IP, the Internet Protocol [RFC791] [RFC2460], is the core protocol of
the Internet. IP defines a simple "connectionless" packet-switched
network. The success of the Internet is largely attributed to IP's
simplicity, the "end-to-end principle" [SRC81] on which the Internet
is based, and the resulting ease of carrying IP on a wide variety of
subnetworks, not necessarily designed with IP in mind. A subnetwork
refers to any network operating immediately below the IP layer to
connect two or more systems using IP (i.e., end hosts or routers).
In its simplest form, this may be a direct connection between the IP
systems (e.g., using a length of cable or a wireless medium).
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This document defines a subnetwork as a layer 2 network, which is a
network that does not rely upon the services of IP routers to forward
packets between parts of the subnetwork. However, IP routers may
bridge frames at Layer 2 between parts of a subnetwork. Sometimes,
it is convenient to aggregate a group of such subnetworks into a
single logical subnetwork. IP routing protocols (e.g., OSPF, IS-IS,
and PIM) can be configured to support this aggregation, but typically
present a layer-3 subnetwork rather than a layer-2 subnetwork. This
may also result in a specific packet passing several times over the
same layer-2 subnetwork via an intermediate layer-3 gateway (router).
Because that aggregation requires layer-3 components, issues thereof
are beyond the scope of this document.
However, while many subnetworks carry IP, they do not necessarily do
so with maximum efficiency, minimum complexity, or cost, nor do they
implement certain features to efficiently support newer Internet
features of increasing importance, such as multicasting or quality of
service.
With the explosive growth of the Internet, IP packets comprise an
increasingly large fraction of the traffic carried by the world's
telecommunications networks. It therefore makes sense to optimize
both existing and new subnetwork technologies for IP as much as
possible.
Optimizing a subnetwork for IP involves three complementary
considerations:
1. Providing functionality sufficient to carry IP.
2. Eliminating unnecessary functions that increase cost or
complexity.
3. Choosing subnetwork parameters that maximize the performance of
the Internet protocols.
Because IP is so simple, consideration 2 is more of an issue than
consideration 1. That is to say, subnetwork designers make many more
errors of commission than errors of omission. However, certain
enhancements to Internet features, such as multicasting and quality-
of-service, benefit significantly from support given by the
underlying subnetworks beyond that necessary to carry "traditional"
unicast, best-effort IP.
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A major consideration in the efficient design of any layered
communication network is the appropriate layer(s) in which to
implement a given function. This issue was first addressed in the
seminal paper, "End-to-End Arguments in System Design" [SRC81]. That
paper argued that many functions can be implemented properly *only*
on an end-to-end basis, i.e., at the highest protocol layers, outside
the subnetwork. These functions include ensuring the reliable
delivery of data and the use of cryptography to provide
confidentiality and message integrity.
Such functions cannot be provided solely by the concatenation of
hop-by-hop services; duplicating these functions at the lower
protocol layers (i.e., within the subnetwork) can be needlessly
redundant or even harmful to cost and performance.
However, partial duplication of functionality in a lower layer can
*sometimes* be justified by performance, security, or availability
considerations. Examples include link-layer retransmission to
improve the performance of an unusually lossy channel, e.g., mobile
radio, link-level encryption intended to thwart traffic analysis, and
redundant transmission links to improve availability, increase
throughput, or to guarantee performance for certain classes of
traffic. Duplication of protocol functions should be done only with
an understanding of system-level implications, including possible
interactions with higher-layer mechanisms.
The original architecture of the Internet was influenced by the
end-to-end principle [SRC81], and has been, in our view, part of the
reason for the Internet's success.
The remainder of this document discusses the various subnetwork
design issues that the authors consider relevant to efficient IP
support.
IPv4 packets (datagrams) vary in size, from 20 bytes (the size of the
IPv4 header alone) to a maximum of 65535 bytes. Subnetworks need not
support maximum-sized (64KB) IP packets, as IP provides a scheme that
breaks packets that are too large for a given subnetwork into
fragments that travel as independent IP packets and are reassembled
at the destination. The maximum packet size supported by a
subnetwork is known as its Maximum Transmission Unit (MTU).
Subnetworks may, but are not required to, indicate the length of each
packet they carry. One example is Ethernet with the widely used DIX
[DIX82] (not IEEE 802.3 [IEEE8023]) header, which lacks a length
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field to indicate the true data length when the packet is padded to a
minimum of 60 bytes. This is not a problem for uncompressed IP
because each IP packet carries its own length field.
If optional header compression [RFC1144] [RFC2507] [RFC2508]
[RFC3095] is used, however, it is required that the link framing
indicate frame length because that is needed for the reconstruction
of the original header.
In IP version 4 (the version now in widespread use), fragmentation
can occur at either the sending host or in an intermediate router,
and fragments can be further fragmented at subsequent routers if
necessary.
In IP version 6 [RFC2460], fragmentation can occur only at the
sending host; it cannot occur in a router (called "router
fragmentation" in this document).
Both IPv4 and IPv6 provide a "path MTU discovery" procedure [RFC1191]
[RFC1435] [RFC1981] that allows the sending host to avoid
fragmentation by discovering the minimum MTU along a given path and
reduce its packet sizes accordingly. This procedure is optional in
IPv4 and IPv6.
Path MTU discovery is widely deployed, but it sometimes encounters
problems. Some routers fail to generate the ICMP messages that
convey path MTU information to the sender, and sometimes the ICMP
messages are blocked by overly restrictive firewalls. The result can
be a "Path MTU Black Hole" [RFC2923] [RFC1435].
The Path MTU Discovery procedure, the persistence of path MTU black
holes, and the deletion of router fragmentation in IPv6 reflect a
consensus of the Internet technical community that router
fragmentation is best avoided. This requires that subnetworks
support MTUs that are "reasonably" large. All IPv4 end hosts are
required to accept and reassemble IP packets of size 576 bytes
[RFC791], but such a small value would clearly be inefficient.
Because IPv6 omits fragmentation by routers, [RFC2460] specifies a
larger minimum MTU of 1280 bytes. Any subnetwork with an internal
packet payload smaller than 1280 bytes must implement a mechanism
that performs fragmentation/reassembly of IP packets to/from
subnetwork frames if it is to support IPv6.
If a subnetwork cannot directly support a "reasonable" MTU with
native framing mechanisms, it should internally fragment. That is,
it should transparently break IP packets into internal data elements
and reassemble them at the other end of the subnetwork.
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This leaves the question of what is a "reasonable" MTU. Ethernet (10
and 100 Mb/s) has an MTU of 1500 bytes, and because of the ubiquity
of Ethernet few Internet paths currently have MTUs larger than this
value. This severely limits the utility of larger MTUs provided by
other subnetworks. Meanwhile, larger MTUs are increasingly desirable
on high-speed subnetworks to reduce the per-packet processing
overhead in host computers, and implementers are encouraged to
provide them even though they may not be usable when Ethernet is also
in the path.
Various "tunneling" schemes, such as GRE [RFC2784] or IP Security in
tunnel mode [RFC2406], treat IP as a subnetwork for IP. Since
tunneling adds header overhead, it can trigger fragmentation, even
when the same physical subnetworks (e.g., Ethernet) are used on both
sides of the host performing IPsec encapsulation. Tunneling has made
it more difficult to avoid router fragmentation and has increased the
incidence of path MTU black holes [RFC2401] [RFC2923]. Larger
subnetwork MTUs may help to alleviate this problem.
In slow networks, the largest possible packet may take a considerable
amount of time to send. This is known as channelisation or
serialisation delay. Total end-to-end interactive response time
should not exceed the well-known human factors limit of 100 to 200
ms. This includes all sources of delay: electromagnetic propagation
delay, queuing delay, serialisation delay, and the store-and-forward
time, i.e., the time to transmit a packet at link speed.
At low link speeds, store-and-forward delays can dominate total
end-to-end delay; these are in turn directly influenced by the
maximum transmission unit (MTU) size. Even when an interactive
packet is given a higher queuing priority, it may have to wait for a
large bulk transfer packet to finish transmission. This worst-case
wait can be set by an appropriate choice of MTU.
For example, if the MTU is set to 1500 bytes, then an MTU-sized
packet will take about 8 milliseconds to send on a T1 (1.536 Mb/s)
link. But if the link speed is 19.2kb/s, then the transmission time
becomes 625 ms -- well above our 100-200ms limit. A 256-byte MTU
would lower this delay to a little over 100 ms. However, care should
be taken not to lower the MTU excessively, as this will increase
header overhead and trigger frequent router fragmentation (if Path
MTU discovery is not in use). This is likely to be the case with
multicast, where Path MTU discovery is ineffective.
One way to limit delay for interactive traffic without imposing a
small MTU is to give priority to this traffic and to preempt (abort)
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the transmission of a lower-priority packet when a higher priority
packet arrives in the queue. However, the link resources used to
send the aborted packet are lost, and overall throughput will
decrease.
Another way to limit delay is to implement a link-level multiplexing
scheme that allows several packets to be in progress simultaneously,
with transmission priority given to segments of higher-priority IP
packets. For links using the Point-To-Point Protocol (PPP)
[RFC1661], multi-class multilink [RFC2686] [RFC2687] [RFC2689]
provides such a facility.
ATM (asynchronous transfer mode), where SNDUs are fragmented and
interleaved across smaller 53-byte ATM cells, is another example of
this technique. However, ATM is generally used on high-speed links
where the store-and-forward delays are already minimal, and it
introduces significant (~9%) increases in overhead due to the
addition of 5-byte cell overhead to each 48-byte ATM cell.
A third example is the Data-Over-Cable Service Interface
Specification (DOCSIS) with typical upstream bandwidths of 2.56 Mb/s
or 5.12 Mb/s. To reduce the impact of a 1500-byte MTU in DOCSIS 1.0
[DOCSIS1], a data link layer fragmentation mechanism is specified in
DOCSIS 1.1 [DOCSIS2]. To accommodate the installed base, DOCSIS 1.1
must be backward compatible with DOCSIS 1.0 cable modems, which
generally do not support fragmentation. Under the co-existence of
DOCSIS 1.0 and DOCSIS 1.1, the unfragmented large data packets from
DOCSIS 1.0 cable modems may affect the quality of service for voice
packets from DOCSIS 1.1 cable modems. In this case, it has been
shown in [DOCSIS3] that the use of bandwidth allocation algorithms
can mitigate this effect.
To summarize, there is a fundamental tradeoff between efficiency and
latency in the design of a subnetwork, and the designer should keep
this tradeoff in mind.
IP requires that subnetworks mark the beginning and end of each
variable-length, asynchronous IP packet. Some examples of links and
subnetworks that do not provide this as an intrinsic feature include:
1. leased lines carrying a synchronous bit stream;
2. ISDN B-channels carrying a synchronous octet stream;
3. dialup telephone modems carrying an asynchronous octet stream;
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and
4. Asynchronous Transfer Mode (ATM) networks carrying an
asynchronous stream of fixed-sized "cells".
The Internet community has defined packet framing methods for all
these subnetworks. The Point-To-Point Protocol (PPP) [RFC1661],
which uses a variant of HDLC, is applicable to bit synchronous,
octet-synchronous, and octet asynchronous links (i.e., examples 1-3
above). PPP is one preferred framing method for IP, since a large
number of systems interoperate with PPP. ATM has its own framing
methods, described in [RFC2684] [RFC2364].
At high speeds, a subnetwork should provide a framed interface
capable of carrying asynchronous, variable-length IP datagrams. The
maximum packet size supported by this interface is discussed above in
the MTU/Fragmentation section. The subnetwork may implement this
facility in any convenient manner.
IP packet boundaries need not coincide with any framing or
synchronization mechanisms internal to the subnetwork. When the
subnetwork implements variable sized data units, the most
straightforward approach is to place exactly one IP packet into each
subnetwork data unit (SNDU), and to rely on the subnetwork's existing
ability to delimit SNDUs to also delimit IP packets. A good example
is Ethernet. However, some subnetworks have SNDUs of one or more
fixed sizes, as dictated by switching, forward error correction
and/or interleaving considerations. Examples of such subnetworks
include ATM, with a single cell payload size of 48 octets plus a 5-
octet header, and IS-95 digital cellular, with two "rate sets" of
four fixed frame sizes each that may be selected on 20 millisecond
boundaries.
Because IP packets are of variable length, they may not necessarily
fit into an integer multiple of fixed-sized SNDUs. An "adaptation
layer" is needed to convert IP packets into SNDUs while marking the
boundary between each IP packet in some manner.
There are several approaches to this problem. The first is to encode
each IP packet into one or more SNDUs with no SNDU containing pieces
of more than one IP packet, and to pad out the last SNDU of the
packet as needed. Bits in a control header added to each SNDU
indicate where the data segment belongs in the IP packet. If the
subnetwork provides in-order, at-most-once delivery, the header can
be as simple as a pair of bits indicating whether the SNDU is the
first and/or the last in the IP packet. Alternatively, for
subnetworks that do not reorder the fragments of an SNDU, only the
last SNDU of the packet could be marked, as this would implicitly
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indicate the next SNDU as the first in a new IP packet. The AAL5
(ATM Adaptation Layer 5) scheme used with ATM is an example of this
approach, though it adds other features, including a payload length
field and a payload CRC.
In AAL5, the ATM User-User Indication, which is encoded in the
Payload Type field of an ATM cell, indicates the last cell of a
packet. The packet trailer is located at the end of the SNDU and
contains the packet length and a CRC.
Another framing technique is to insert per-segment overhead to
indicate the presence of a segment option. When present, the option
carries a pointer to the end of the packet. This differs from AAL5
in that it permits another packet to follow within the same segment.
MPEG-2 Transport Streams [EN301192] [ISO13818] support this style of
fragmentation, and may either use padding (limiting each MPEG
transport stream packet to carry only part of one IP packet), or
allow a second IP packet to start in the same Transport Stream packet
(no padding).
A third approach is to insert a special flag sequence into the data
stream between each IP packet, and to pack the resulting data stream
into SNDUs without regard to SNDU boundaries. This may have
implications when frames are lost. The flag sequence can also pad
unused space at the end of an SNDU. If the special flag appears in
the user data, it is escaped to an alternate sequence (usually larger
than a flag) to avoid being misinterpreted as a flag. The HDLC-based
framing schemes used in PPP are all examples of this approach.
All three adaptation schemes introduce overhead; how much depends on
the distribution of IP packet sizes, the size(s) of the SNDUs, and in
the HDLC-like approaches, the content of the IP packet (since flag-
like sequences occurring in the packet must be escaped, which expands
them). The designer must also weigh implementation complexity and
performance in the choice and design of an adaptation layer.
IP has no notion of a "connection"; it is a purely connectionless
protocol. When a connection is required by an application, it is
usually provided by TCP [RFC793], the Transmission Control Protocol,
running atop IP on an end-to-end basis.
Connection-oriented subnetworks can be (and are widely) used to carry
IP, but often with considerable complexity. Subnetworks consisting
of few nodes can simply open a permanent connection between each pair
of nodes. This is frequently done with ATM. However, the number of
connections increases as the square of the number of nodes, so this
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is clearly impractical for large subnetworks. A "shim" layer between
IP and the subnetwork is therefore required to manage connections.
This is one of the most common functions of a Subnetwork Dependent
Convergence Function (SNDCF) sublayer between IP and a subnetwork.
SNDCFs typically open subnetwork connections as needed when an IP
packet is queued for transmission and close them after an idle
timeout. There is no relation between subnetwork connections and any
connections that may exist at higher layers (e.g., TCP).
Because Internet traffic is typically bursty and transaction-
oriented, it is often difficult to pick an optimal idle timeout. If
the timeout is too short, subnetwork connections are opened and
closed rapidly, possibly over-stressing the subnetwork connection
management system (especially if it was designed for voice traffic
call holding times). If the timeout is too long, subnetwork
connections are idle much of the time, wasting any resources
dedicated to them by the subnetwork.
Purely connectionless subnets (such as Ethernet), which have no state
and dynamically share resources, are optimal for supporting best-
effort IP, which is stateless and dynamically shares resources.
Connection-oriented packet networks (such as ATM and Frame Relay),
which have state and dynamically share resources, are less optimal,
since best-effort IP does not benefit from the overhead of creating
and maintaining state. Connection-oriented circuit-switched networks
(including the PSTN and ISDN) have state and statically allocate
resources for a call, and thus require state creation and maintenance
overhead, but do not benefit from the efficiencies of statistical
multiplexing sharing of capacity inherent in IP.
In any event, if an SNDCF that opens and closes subnet connections is
used to support IP, care should be taken to make sure that connection
processing in the subnet can keep up with relatively short holding
times.
Subnetworks fall into two categories: point-to-point and shared. A
point-to-point subnet has exactly two endpoint components (hosts or
routers); a shared link has more than two endpoint components, using
either an inherently broadcast medium (e.g., Ethernet, radio) or a
switching layer hidden from the network layer (e.g., switched
Ethernet, Myrinet [MYR95], ATM). Switched subnetworks handle
broadcast by copying broadcast packets, providing each interface that
supports one, or more, systems (hosts or routers) with a copy of each
packet.
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Several Internet protocols for IPv4 make use of broadcast
capabilities, including link-layer address lookup (ARP), auto-
configuration (RARP, BOOTP, DHCP), and routing (RIP).
A lack of broadcast capability can impede the performance of these
protocols, or render them inoperable (e.g., DHCP). ARP-like link
address lookup can be provided by a centralized database, but at the
expense of potentially higher response latency and the need for nodes
to have explicit knowledge of the ARP server address. Shared links
should support native, link-layer subnet broadcast.
A corresponding set of IPv6 protocols uses multicasting (see next
section) instead of broadcasting to provide similar functions with
improved scaling in large networks.
The Internet model includes "multicasting", where IP packets are sent
to all the members of a multicast group [RFC1112] [RFC3376]
[RFC2710]. Multicast is an option in IPv4, but a standard feature of
IPv6. IPv4 multicast is currently used by multimedia,
teleconferencing, gaming, and file distribution (web, peer-to-peer
sharing) applications, as well as by some key network and host
protocols (e.g., RIPv2, OSPF, NTP). IPv6 additionally relies on
multicast for network configuration (DHCP-like autoconfiguration) and
link-layer address discovery [RFC2461] (replacing ARP). In the case
of IPv6, this can allow autoconfiguration and address discovery to
span across routers, whereas the IPv4 broadcast-based services cannot
without ad-hoc router support [RFC1812].
Multicast-enabled IP routers organize each multicast group into a
spanning tree, and route multicast packets by making copies of each
multicast packet and forwarding the copies to each output interface
that includes at least one downstream member of the multicast group.
Multicasting is considerably more efficient when a subnetwork
explicitly supports it. For example, a router relaying a multicast
packet onto an Ethernet segment need send only one copy of the
packet, no matter how many members of the multicast group are
connected to the segment. Without native multicast support, routers
and switches on shared links would need to use broadcast with
software filters, such that every multicast packet sent incurs
software overhead for every node on the subnetwork, even if a node is
not a member of the multicast group. Alternately, the router would
transmit a separate copy to every member of the multicast group on
the segment, as is done on multicast-incapable switched subnets.
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Subnetworks using shared channels (e.g., radio LANs, Ethernets) are
especially suitable for native multicasting, and their designers
should make every effort to support it. This involves designating a
section of the subnetwork's own address space for multicasting. On
these networks, multicast is basically broadcast on the medium, with
Layer-2 receiver filters.
Subnet interfaces also need to be designed to accept packets
addressed to some number of multicast addresses, in addition to the
unicast packets specifically addressed to them. The number of
multicast addresses that needs to be supported by a host depends on
the requirements of the associated host; at least several dozen will
meet most current needs.
On low-speed networks, the multicast address recognition function may
be readily implemented in host software, but on high-speed networks,
it should be implemented in subnetwork hardware. This hardware need
not be complete; for example, many Ethernet interfaces implement a
"hashing" function where the IP layer receives all of the multicast
(and unicast) traffic to which the associated host subscribes, plus
some small fraction of multicast traffic to which the host does not
subscribe. Host/router software then has to discard the unwanted
packets that pass the Layer-2 multicast address filter [RFC1112].
There does not need to be a one-to-one mapping between a Layer-2
multicast address and an IP multicast address. An address overlap
may significantly degrade the filtering capability of a receiver's
hardware multicast address filter. A subnetwork supporting only
broadcast should use this service for multicast and must rely on
software filtering.
Switched subnetworks must also provide a mechanism for copying
multicast packets to ensure the packets reach at least all members of
a multicast group. One option is to "flood" multicast packets in the
same manner as broadcast. This can lead to unnecessary transmissions
on some subnetwork links (notably non-multicast-aware Ethernet
switches). Some subnetworks therefore allow multicast filter tables
to control which links receive packets belonging to a specific group.
To configure this automatically requires access to Layer-3 group
membership information (e.g., IGMP [RFC3376], or MLD [RFC2710]).
Various implementation options currently exist to provide a subnet
node with a list of mappings of multicast addresses to
ports/interfaces. These employ a range of approaches, including
signaling from end hosts (e.g., IEEE 802 GARP/GMRP [802.1p]),
signaling from switches (e.g., CGMP [CGMP] and RGMP [RFC3488]),
interception and proxy of IP group membership packets (e.g., IGMP/MLD
Proxy [MAGMA-PROXY]), and enabling Layer-2 devices to
snoop/inspect/peek into forwarded Layer-3 protocol headers (e.g.,
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RFC 3819 Advice for Internet Subnetwork Designers July 2004
IGMP, MLD, PIM) so that they may infer Layer-3 multicast group
membership [MAGMA-SNOOP]. These approaches differ in their
complexity, flexibility, and ability to support new protocols.
Some subnets allow a number of subnet nodes to share a channel
efficiently by assigning transmission opportunities dynamically.
Transmission opportunities are requested by a subnet node when it has
packets to send. The subnet schedules and grants transmission
opportunities sufficient to allow the transmitting subnet node to
send one or more packets (or packet fragments). We call these
subnets Bandwidth on Demand (BoD) subnets. Examples of BoD subnets
include Demand Assignment Multiple Access (DAMA) satellite and
terrestrial wireless networks, IEEE 802.11 point coordination
function (PCF) mode, and DOCSIS. A connection-oriented network (such
as the PSTN, ATM or Frame Relay) reserves resources on a much longer
timescale, and is therefore not a BoD subnet in our taxonomy.
The design parameters for BoD are similar to those in connection-
oriented subnetworks, although the implementations may vary
significantly. In BoD, the user typically requests access to the
shared channel for some duration. Access may be allocated for a
period of time at a specific rate, for a certain number of packets,
or until the user releases the channel. Access may be coordinated
through a central management entity or with a distributed algorithm
amongst the users. Examples of the resource that may be shared
include a terrestrial wireless hop, an upstream channel in a cable
television system, a satellite uplink, and an end-to-end satellite
channel.
Long-delay BoD subnets pose problems similar to connection-oriented
subnets in anticipating traffic. While connection-oriented subnets
hold idle channels open expecting new data to arrive, BoD subnets
request channel access based on buffer occupancy (or expected buffer
occupancy) on the sending port. Poor performance will likely result
if the sender does not anticipate additional traffic arriving at that
port during the time it takes to grant a transmission request. It is
recommended that the algorithm have the capability to extend a hold
on the channel for data that has arrived after the original request
was generated (this may be done by piggybacking new requests on user
data).
There is a wide variety of BoD protocols available. However, there
has been relatively little comprehensive research on the interactions
between BoD mechanisms and Internet protocol performance. Research
on some specific mechanisms is available (e.g., [AR02]). One item
that has been studied is TCP's retransmission timer [KY02]. BoD
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systems can cause spurious timeouts when adjusting from a relatively
high data rate, to a relatively low data rate. In this case, TCP's
transmitted data takes longer to get through the network than
predicted by the TCP sender's computed retransmission timeout.
Therefore, the TCP sender is prone to resending a segment
prematurely.
In the Internet architecture, the ultimate responsibility for error
recovery is at the end points [SRC81]. The Internet may occasionally
drop, corrupt, duplicate, or reorder packets, and the transport
protocol (e.g., TCP) or application (e.g., if UDP is used as the
transport protocol) must recover from these errors on an end-to-end
basis [RFC3155]. Error recovery in the subnetwork is therefore
justifiable only to the extent that it can enhance overall
performance. It is important to recognize that a subnetwork can go
too far in attempting to provide error recovery services in the
Internet environment. Subnet reliability should be "lightweight",
i.e., it only has to be "good enough", *not* perfect.
In this section, we discuss how to analyze characteristics of a
subnetwork to determine what is "good enough". The discussion below
focuses on TCP, which is the most widely-used transport protocol in
the Internet. It is widely believed (and is a stated goal within the
IETF) that non-TCP transport protocols should attempt to be "TCP-
friendly" and have many of the same performance characteristics.
Thus, the discussion below should be applicable, even to portions of
the Internet where TCP may not be the predominant protocol.
Error recovery involves the generation and transmission of redundant
information computed from user data. Depending on how much redundant
information is sent and how it is generated, the receiver can use it
to reliably detect transmission errors, correct up to some maximum
number of transmission errors, or both. The general approach is
known as Error Control Coding, or ECC.
The use of ECC to detect transmission errors so that retransmissions
(hopefully without errors) can be requested is widely known as "ARQ"
(Automatic Repeat Request).
When enough ECC information is available to permit the receiver to
correct some transmission errors without a retransmission, the
approach is known as Forward Error Correction (FEC). Due to the
greater complexity of the required ECC and the need to tailor its
design to the characteristics of a specific modem and channel, FEC
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has traditionally been implemented in special-purpose hardware
integral to a modem. This effectively makes it part of the physical
layer.
Unlike ARQ, FEC was rarely used for telecommunications outside of
space links prior to the 1990s. It is now nearly universal in
telephone, cable and DSL modems, digital satellite links, and digital
mobile telephones. FEC is also heavily used in optical and magnetic
storage where "retransmissions" are not possible.
Some systems use hybrid combinations of ARQ layered atop FEC; V.90
dialup modems (in the upstream direction) with V.42 error control are
one example. Most errors are corrected by the trellis (FEC) code
within the V.90 modem, and most remaining errors are detected and
corrected by the ARQ mechanisms in V.42.
Work is now underway to apply FEC above the physical layer, primarily
in connection with reliable multicasting [RFC3048] [RFC3450-RFC3453]
where conventional ARQ mechanisms are inefficient or difficult to
implement. However, in this discussion, we will assume that if FEC
is present, it is implemented within the physical layer.
Depending on the layer in which it is implemented, error control can
operate on an end-to-end basis or over a shorter span, such as a
single link. TCP is the most important example of an end-to-end
protocol that uses an ARQ strategy.
Many link-layer protocols use ARQ, usually some flavor of HDLC
[ISO3309]. Examples include the X.25 link layer, the AX.25 protocol
used in amateur packet radio, 802.11 wireless LANs, and the reliable
link layer specified in IEEE 802.2.
Only end-to-end error recovery can ensure reliable service to the
application (see Section 8). However, some subnetworks (e.g., many
wireless links) also have link-layer error recovery as a performance
enhancement [RFC3366]. For example, many cellular links have small
physical frame sizes (< 100 bytes) and relatively high frame loss
rates. Relying solely on end-to-end error recovery can clearly yield
a performance degradation, as retransmissions across the end-to-end
path take much longer to be received than when link layer
retransmissions are used. Thus, link-layer error recovery can often
increase end-to-end performance. As a result, link-layer and end-
to-end recovery often co-exist; this can lead to the possibility of
inefficient interactions between the two layers of ARQ protocols.
This inter-layer "competition" might lead to the following wasteful
situation. When the link layer retransmits (parts of) a packet, the
link latency momentarily increases. Since TCP bases its
Karn, et al. Best Current Practice [Page 15]
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retransmission timeout on prior measurements of total end-to-end
latency, including that of the link in question, this sudden increase
in latency may trigger an unnecessary retransmission by TCP of a
packet that the link layer is still retransmitting. Such spurious
end-to-end retransmissions generate unnecessary load and reduce end-
to-end throughput. As a result, the link layer may even have
multiple copies of the same packet in the same link queue at the same
time. In general, one could say the competing error recovery is
caused by an inner control loop (link-layer error recovery) reacting
to the same signal as an outer control loop (end-to-end error
recovery) without any coordination between the loops. Note that this
is solely an efficiency issue; TCP continues to provide reliable
end-to-end delivery over such links.
This raises the question of how persistent a link-layer sender should
be in performing retransmission [RFC3366]. We define the link-layer
(LL) ARQ persistency as the maximum time that a particular link will
spend trying to transfer a packet before it can be discarded. This
deliberately simplified definition says nothing about the maximum
number of retransmissions, retransmission strategies, queue sizes,
queuing disciplines, transmission delays, or the like. The reason we
use the term LL ARQ persistency, instead of a term such as "maximum
link-layer packet holding time," is that the definition closely
relates to link-layer error recovery. For example, on links that
implement straightforward error recovery strategies, LL ARQ
persistency will often correspond to a maximum number of
retransmissions permitted per link-layer frame.
For link layers that do not or cannot differentiate between flows
(e.g., due to network layer encryption), the LL ARQ persistency
should be small. This avoids any harmful effects or performance
degradation resulting from indiscriminate high persistence. A
detailed discussion of these issues is provided in [RFC3366].
However, when a link layer can identify individual flows and apply
ARQ selectively [LKJK02], then the link ARQ persistency should be
high for a flow using reliable unicast transport protocols (e.g.,
TCP) and must be low for all other flows. Setting the link ARQ
persistency larger than the largest link outage allows TCP to rapidly
restore transmission without needing to wait for a retransmission
time out. This generally improves TCP performance in the face of
transient outages. However, excessively high persistence may be
disadvantageous; a practical upper limit of 30-60 seconds may be
desirable. Implementation of such schemes remains a research issue.
(See also the following section "Recovery from Subnetwork Outages").
Karn, et al. Best Current Practice [Page 16]
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Many subnetwork designers have opportunities to reduce the
probability of packet loss, e.g., with FEC, ARQ, and interleaving, at
the cost of increased delay. TCP performance improves with
decreasing loss but worsens with increasing end-to-end delay, so it
is important to find the proper balance through analysis and
simulation.
Some types of subnetworks, particularly mobile radio, are subject to
frequent temporary outages. For example, an active cellular data
user may drive or walk into an area (such as a tunnel) that is out of
range of any base station. No packets will be delivered successfully
until the user returns to an area with coverage.
The Internet protocols currently provide no standard way for a
subnetwork to explicitly notify an upper layer protocol (e.g., TCP)
that it is experiencing an outage rather than severe congestion.
Under these circumstances TCP will, after each unsuccessful
retransmission, wait even longer before trying again; this is its
"exponential back-off" algorithm. Furthermore, TCP will not discover
that the subnetwork outage has ended until its next retransmission
attempt. If TCP has backed off, this may take some time. This can
lead to extremely poor TCP performance over such subnetworks.
It is therefore highly desirable that a subnetwork subject to outages
does not silently discard packets during an outage. Ideally, the
subnetwork should define an interface to the next higher layer (i.e.,
IP) that allows it to refuse packets during an outage, and to
automatically ask IP for new packets when it is again able to deliver
them. If it cannot do this, then the subnetwork should hold onto at
least some of the packets it accepts during an outage and attempt to
deliver them when the outage ends. When packets are discarded, IP
should be notified so that the appropriate ICMP messages can be sent.
Note that it is *not* necessary to completely avoid dropping packets
during an outage. The purpose of holding onto a packet during an
outage, either in the subnetwork or at the IP layer, is so that its
eventual delivery will implicitly notify TCP that the subnetwork is
again operational. This is to enhance performance, not to ensure
reliability -- reliability, as discussed earlier, can only be ensured
on an end-to-end basis.
Only a few packets per TCP connection, including ACKs, need be held
in this way to cause the TCP sender to recover from the additional
losses once the flow resumes [RFC3366].
Karn, et al. Best Current Practice [Page 17]
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Because it would be a layering violation (and possibly a performance
hit) for IP or a subnetwork layer to look at TCP headers (which would
in any event be impossible if IPsec encryption [RFC2401] is in use),
it would be reasonable for the IP or subnetwork layers to choose, as
a design parameter, some small number of packets that will be
retained during an outage.
The TCP [RFC793], UDP [RFC768], ICMP, and IPv4 [RFC791] protocols all
use the same simple 16-bit 1's complement checksum algorithm
[RFC1071] to detect corrupted packets. The IPv4 header checksum
protects only the IPv4 header, while the TCP, ICMP, and UDP checksums
provide end-to-end error detection for both the transport pseudo
header (including network and transport layer information) and the
transport payload data. Protection of the data is optional for
applications using UDP [RFC768] for IPv4, but is required for IPv6.
The Internet checksum is not very strong from a coding theory
standpoint, but it is easy to compute in software, and various
proposals to replace the Internet checksums with stronger checksums
have failed. However, it is known that undetected errors can and do
occur in packets received by end hosts [SP2000].
To reduce processing costs, IPv6 has no IP header checksum. The
destination host detects "important" errors in the IP header, such as
the delivery of the packet to the wrong destination. This is done by
including the IP source and destination addresses (pseudo header) in
the computation of the checksum in the TCP or UDP header, a practice
already performed in IPv4. Errors in other IPv6 header fields may go
undetected within the network; this was considered a reasonable price
to pay for a considerable reduction in the processing required by
each router, and it was assumed that subnetworks would use a strong
link CRC.
One way to provide additional protection for an IPv4 or IPv6 header
is by the authentication and packet integrity services of the IP
Security (IPsec) protocol [RFC2401]. However, this may not be a
choice available to the subnetwork designer.
Most subnetworks implement error detection just above the physical
layer. Packets corrupted in transmission are detected and discarded
before delivery to the IP layer. A 16-bit cyclic redundancy check
(CRC) is usually the minimum for error detection. This is
significantly more robust against most patterns of errors than the
16-bit Internet checksum. Note that the error detection properties
of a specific CRC code diminish with increasing frame size. The
Point-to-Point Protocol [RFC1662] requires support of a 16-bit CRC
Karn, et al. Best Current Practice [Page 18]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
for each link frame, with a 32-bit CRC as an option. (PPP is often
used in conjunction with a dialup modem, which provides its own error
control). Other subnetworks, including 802.3/Ethernet, AAL5/ATM,
FDDI, Token Ring, and PPP over SONET/SDH all use a 32-bit CRC. Many
subnetworks can also use other mechanisms to enhance the error
detection capability of the link CRC (e.g., FEC in dialup modems,
mobile radio and satellite channels).
Any new subnetwork designed to carry IP should therefore provide
error detection for each IP packet that is at least as strong as the
32-bit CRC specified in [ISO3309]. While this will achieve a very
low undetected packet error rate due to transmission errors, it will
not (and need not) achieve a very low packet loss rate as the
Internet protocols are better suited to dealing with lost packets
than to dealing with corrupted packets [SRC81].
Packet corruption may be, and is, also caused by bugs in host and
router hardware and software. Even if every subnetwork implemented
strong error detection, it is still essential that end-to-end
checksums are used at the receiving end host [SP2000].
Designers of complex subnetworks consisting of internal links and
packet switches should consider implementing error detection on an
edge-to-edge basis to cover an entire SNDU (or IP packet). A CRC
would be generated at the entry point to the subnetwork and checked
at the exit endpoint. This may be used instead of, or in combination
with, error detection at the interface to each physical link. An
edge-to-edge check has the significant advantage of protecting
against errors introduced anywhere within the subnetwork, not just
within its transmission links. Examples of this approach include the
way in which the Ethernet CRC-32 is handled by LAN bridges [802.1D].
ATM AAL5 [ITU-I363] also uses an edge-to-edge CRC-32.
Some specific applications may be tolerant of residual errors in the
data they exchange, but removal of the link CRC may expose the
network to an undesirable increase in undetected errors in the IP and
transport headers. Applications may also require a high level of
error protection for control information exchanged by protocols
acting above the transport layer. One example is a voice codec,
which is robust against bit errors in the speech samples. For such
mechanisms to work, the receiving application must be able to
tolerate receiving corrupted data. This also requires that an
application uses a mechanism to signal that payload corruption is
permitted and to indicate the coverage (headers and data) required to
be protected by the subnetwork CRC. The UDP-Lite protocol [RFC3828]
is the first Internet standards track transport protocol supporting
partial payload protection. Receipt of corrupt data by arbitrary
Karn, et al. Best Current Practice [Page 19]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
application protocols carries a serious danger that a subnet delivers
data with errors that remain undetected by the application and hence
corrupt the communicated data [SRC81].
One of TCP's functions is end-host based congestion control for the
Internet. This is a critical part of the overall stability of the
Internet, so it is important that link-layer designers understand
TCP's congestion control algorithms.
TCP assumes that, at the most abstract level, the network consists of
links and queues. Queues provide output-buffering on links that are
momentarily oversubscribed. They smooth instantaneous traffic bursts
to fit the link bandwidth. When demand exceeds link capacity long
enough to fill the queue, packets must be dropped. The traditional
action of dropping the most recent packet ("tail dropping") is no
longer recommended [RFC2309] [RFC2914], but it is still widely
practiced.
TCP uses sequence numbering and acknowledgments (ACKs) on an
end-to-end basis to provide reliable, sequenced delivery. TCP ACKs
are cumulative, i.e., each implicitly ACKs every segment received so
far. If a packet with an unexpected sequence number is received, the
ACK field in the packets returned by the receiver will cease to
advance. Using an optional enhancement, TCP can send selective
acknowledgments (SACKs) [RFC2018] to indicate which segments have
arrived at the receiver.
Since the most common cause of packet loss is congestion, TCP treats
packet loss as an indication of potential Internet congestion along
the path between TCP end hosts. This happens automatically, and the
subnetwork need not know anything about IP or TCP. A subnetwork node
simply drops packets whenever it must, though some packet-dropping
strategies (e.g., RED) are more fair to competing flows than others.
TCP recovers from packet losses in two different ways. The most
important mechanism is the retransmission timeout. If an ACK fails
to arrive after a certain period of time, TCP retransmits the oldest
unacked packet. Taking this as a hint that the network is congested,
TCP waits for the retransmission to be ACKed before it continues, and
it gradually increases the number of packets in flight as long as a
timeout does not occur again.
A retransmission timeout can impose a significant performance
penalty, as the sender is idle during the timeout interval and
restarts with a congestion window of one TCP segment following the
Karn, et al. Best Current Practice [Page 20]
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timeout. To allow faster recovery from the occasional lost packet in
a bulk transfer, an alternate scheme, known as "fast recovery", was
introduced [RFC2581] [RFC2582] [RFC2914] [TCPF98].
Fast recovery relies on the fact that when a single packet is lost in
a bulk transfer, the receiver continues to return ACKs to subsequent
data packets that do not actually acknowledge any newly-received
data. These are known as "duplicate acknowledgments" or "dupacks".
The sending TCP can use dupacks as a hint that a packet has been lost
and retransmit it without waiting for a timeout. Dupacks effectively
constitute a negative acknowledgment (NAK) for the packet sequence
number in the acknowledgment field. TCP waits until a certain number
of dupacks (currently 3) are seen prior to assuming a loss has
occurred; this helps avoid an unnecessary retransmission during
out-of-sequence delivery.
A technique called "Explicit Congestion Notification" (ECN) [RFC3168]
allows routers to directly signal congestion to hosts without
dropping packets. This is done by setting a bit in the IP header.
Since ECN support is likely to remain optional, the lack of an ECN
bit must *never* be interpreted as a lack of congestion. Thus, for
the foreseeable future, TCP must interpret a lost packet as a signal
of congestion.
The TCP "congestion avoidance" [RFC2581] algorithm maintains a
congestion window (cwnd) controlling the amount of data TCP may have
in flight at any moment. Reducing cwnd reduces the overall bandwidth
obtained by the connection; similarly, raising cwnd increases
performance, up to the limit of the available capacity.
TCP probes for available network capacity by initially setting cwnd
to one or two packets and then increasing cwnd by one packet for each
ACK returned from the receiver. This is TCP's "slow start"
mechanism. When a packet loss is detected (or congestion is signaled
by other mechanisms), cwnd is reset to one and the slow start process
is repeated until cwnd reaches one half of its previous setting
before the reset. Cwnd continues to increase past this point, but at
a much slower rate than before. If no further losses occur, cwnd
will ultimately reach the window size advertised by the receiver.
This is an "Additive Increase, Multiplicative Decrease" (AIMD)
algorithm. The steep decrease of cwnd in response to congestion
provides for network stability; the AIMD algorithm also provides for
fairness between long running TCP connections sharing the same path.
Karn, et al. Best Current Practice [Page 21]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
Caveat
Here we present a current "state-of-the-art" understanding of TCP
performance. This analysis attempts to characterize the performance
of TCP connections over links of varying characteristics.
Link designers may wish to use the techniques in this section to
predict what performance TCP/IP may achieve over a new link-layer
design. Such analysis is encouraged. Because this is a relatively
new analysis, and the theory is based on single-stream TCP
connections under "ideal" conditions, it should be recognized that
the results of such analysis may differ from actual performance in
the Internet. That being said, we have done our best to provide the
designers with helpful information to get an accurate picture of the
capabilities and limitations of TCP under various conditions.
The performance of TCP's AIMD Congestion Avoidance algorithm has been
extensively analyzed. The current best formula for the performance
of the specific algorithms used by Reno TCP (i.e., the TCP specified
in [RFC2581]) is given by Padhye, et al. [PFTK98]. This formula is:
MSS
BW = --------------------------------------------------------
RTT*sqrt(1.33*p) + RTO*p*[1+32*p^2]*min[1,3*sqrt(.75*p)]
where
BW is the maximum TCP throughout achievable by an
individual TCP flow
MSS is the TCP segment size being used by the connection
RTT is the end-to-end round trip time of the TCP connection
RTO is the packet timeout (based on RTT)
p is the packet loss rate for the path
(i.e., .01 if there is 1% packet loss)
Note that the speed of the links making up the Internet path does not
explicitly appear in this formula. Attempting to send faster than
the slowest link in the path causes the queue to grow at the
transmitter driving the bottleneck. This increases the RTT, which in
turn reduces the achievable throughput.
This is currently considered to be the best approximate formula for
Reno TCP performance. A further simplification of this formula is
generally made by assuming that RTO is approximately 5*RTT.
Karn, et al. Best Current Practice [Page 22]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
TCP is constantly being improved. A simpler formula, which gives an
upper bound on the performance of any AIMD algorithm which is likely
to be implemented in TCP in the future, was derived by Ott, et al.
[MSMO97].
MSS 1
BW = C --- -------
RTT sqrt(p)
where C is 0.93.
Both formulae assume that the TCP Receiver Window is not limiting the
performance of the connection. Because the receiver window is
entirely determined by end-hosts, we assume that hosts will maximize
the announced receiver window to maximize their network performance.
Both of these formulae allow BW to become infinite if there is no
loss. However, an Internet path will drop packets at bottlenecked
queues if the load is too high. Thus, a completely lossless TCP/IP
network can never occur (unless the network is being underutilized).
The RTT used is the arithmetic average, including queuing delays.
The formulae are for a single TCP connection. If a path carries many
TCP connections, each will follow the formulae above independently.
The formulae assume long-running TCP connections. For connections
that are extremely short (<10 packets) and don't lose any packets,
performance is driven by the TCP slow-start algorithm. For
connections of medium length, where on average only a few segments
are lost, single connection performance will actually be slightly
better than given by the formulae above.
The difference between the simple and complex formulae above is that
the complex formula includes the effects of TCP retransmission
timeouts. For very low levels of packet loss (significantly less
than 1%), timeouts are unlikely to occur, and the formulae lead to
very similar results. At higher packet losses (1% and above), the
complex formula gives a more accurate estimate of performance (which
will always be significantly lower than the result from the simple
formula).
Note that these formulae break down as p approaches 100%.
Karn, et al. Best Current Practice [Page 23]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
Consider the following example:
A designer invents a new wireless link layer which, on average, loses
1% of IP packets. The link layer supports packets of up to 1040
bytes, and has a one-way delay of 20 msec.
If this link were to be used on an Internet path with a round trip
time greater than 80ms, the upper bound may be computed by:
For MSS, use 1000 bytes to exclude the 40 bytes of minimum IPv4 and
TCP headers.
For RTT, use 120 msec (80 msec for the Internet part, plus 20 msec
each way for the new wireless link).
For p, use .01. For C, assume 1.
The simple formula gives:
BW = (1000 * 8 bits) / (.120 sec * sqrt(.01)) = 666 kbit/sec
The more complex formula gives:
BW = 402.9 kbit/sec
If this were a 2 Mb/s wireless LAN, the designers might be somewhat
disappointed.
Some observations on performance:
1. We have assumed that the packet losses on the link layer are
interpreted as congestion by TCP. This is a "fact of life" that
must be accepted.
2. The equations for TCP performance are all expressed in terms of
packet loss, but many subnetwork designers think in terms of
bit-error ratio. *If* channel bit errors are independent, then
the probability of a packet being corrupted is:
p = 1 - ([1 - BER]^[FRAME_SIZE*8])
Here we assume FRAME_SIZE is in bytes and "^" represents
exponentiation. It includes the user data and all headers
(TCP,IP and subnetwork). (Note: this analysis assumes the
Karn, et al. Best Current Practice [Page 24]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
subnetwork does not perform ARQ or transparent fragmentation
[RFC3366].) If the inequality
BER * [FRAME_SIZE*8] << 1
holds, the packet loss probability p can be approximated by:
p = BER * [FRAME_SIZE*8]
These equations can be used to apply BER to the performance
equations above.
Note that FRAME_SIZE can vary from one packet to the next. Small
packets (such as TCP acks) generally have a smaller probability
of packet error than, say, a TCP packet carrying one MSS (maximum
segment size) of user data. A flow of small TCP acks can be
expected to be slightly more reliable than a stream of larger TCP
data segments.
It bears repeating that the above analysis assumes that bit
errors are statistically independent. Because this is not true
for many real links, our computation of p is actually an upper
bound, not the exact probability of packet loss.
There are many reasons why bit errors are not independent on real
links. Many radio links are affected by propagation fading or by
interference that lasts over many bit times. Also, links with
Forward Error Correction (FEC) generally have very non-uniform
bit error distributions that depend on the type of FEC, but in
general the uncorrected errors tend to occur in bursts even when
channel symbol errors are independent. In all such cases, our
computation of p from BER can only place an upper limit on the
packet loss rate.
If the distribution of errors under the FEC scheme is known, one
could apply the same type of analysis as above, using the correct
distribution function for the BER. It is more likely in these
FEC cases, however, that empirical methods are needed to
determine the actual packet loss rate.
3. Note that the packet size plays an important role. If the
subnetwork loss characteristics are such that large packets have
the same probability of loss as smaller packets, then larger
packets will yield improved performance.
Karn, et al. Best Current Practice [Page 25]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
4. We have chosen a specific RTT that might occur on a wide-area
Internet path within the USA. It is important to recognize that
a variety of RTT values are experienced in the Internet.
For example, RTTs are typically less than 10 msec in a wired LAN
environment when communicating with a local host. International
connections may have RTTs of 200 msec or more. Modems and other
low-capacity links can add considerable delay due to their long
packet transmission (serialisation) times.
Links over geostationary repeater satellites have one-way speed-
of-light delays of around 250ms, a minimum of 125ms propagation
delay up to the satellite and 125ms down. The RTT of an end-to-
end TCP connection that includes such a link can be expected to
be greater than 250ms.
Queues on heavily-congested links may back up, increasing RTTs.
Finally, virtual private networks (VPNs) and other forms of
encryption and tunneling can add significant end-to-end delay to
network connections.
It is generally recognized that specific service guarantees are
needed to support real-time multimedia, toll-quality telephony, and
other performance-critical applications. The provision of such
Quality of Service guarantees in the Internet is an active area of
research and standardization. The IETF has not converged on a single
service model, set of services, or single mechanism that will offer
useful guarantees to applications and be scalable to the Internet.
Indeed, the IETF does not have a single definition of Quality of
Service. [RFC2990] represents a current understanding of the
challenges in architecting QoS for the Internet.
There are presently two architectural approaches to providing
mechanisms for QoS support in the Internet.
IP Integrated Services (Intserv) [RFC1633] provides fine-grained
service guarantees to individual flows. Flows are identified by a
flow specification (flowspec), which creates a stateful association
between individual packets by matching fields in the packet header.
Capacity is reserved for the flow, and appropriate traffic
conditioning and scheduling is installed in routers along the path.
The ReSerVation Protocol (RSVP) [RFC2205] [RFC2210] is usually, but
need not necessarily be, used to install the flow QoS state. Intserv
defines two services, in addition to the Default (best effort)
service.
Karn, et al. Best Current Practice [Page 26]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
1. Guaranteed Service (GS) [RFC2212] offers hard upper bounds on
delay to flows that conform to a traffic specification (TSpec).
It uses a fluid-flow model to relate the TSpec and reserved
bandwidth (RSpec) to variable delay. Non-conforming packets are
forwarded on a best-effort basis.
2. Controlled Load Service (CLS) [RFC2211] offers delay and packet
loss equivalent to that of an unloaded network to flows that
conform to a TSpec, but no hard bounds. Non-conforming packets
are forwarded on a best-effort basis.
Intserv requires installation of state information in every
participating router. Performance guarantees cannot be made unless
this state is present in every router along the path. This, along
with RSVP processing and the need for usage-based accounting, is
believed to have scalability problems, particularly in the core of
the Internet [RFC2208].
IP Differentiated Services (Diffserv) [RFC2475] provides a "toolkit"
offering coarse-grained controls to aggregates of flows. Diffserv in
itself does *not* provide QoS guarantees, but can be used to
construct services with QoS guarantees across a Diffserv domain.
Diffserv attempts to address the scaling issues associated with
Intserv by requiring state awareness only at the edge of a Diffserv
domain. At the edge, packets are classified into flows, and the
flows are conditioned (marked, policed, or shaped) to a traffic
conditioning specification (TCS). A Diffserv Codepoint (DSCP),
identifying a per-hop behavior (PHB), is set in each packet header.
The DSCP is carried in the DS-field, subsuming six bits of the former
Type-of-Service (ToS) byte [RFC791] of the IP header [RFC2474]. The
PHB denotes the forwarding behavior to be applied to the packet in
each node in the Diffserv domain. Although there is a "recommended"
DSCP associated with each PHB, the mappings from DSCPs to PHBs are
defined by the DS-domain. In fact, there can be several DSCPs
associated with the same PHB. Diffserv presently defines three PHBs.
1. The class selector PHB [RFC2474] replaces the IP precedence field
of the former ToS byte. It offers relative forwarding
priorities.
2. The Expedited Forwarding (EF) PHB [RFC3246] [RFC3248] guarantees
that packets will have a well-defined minimum departure rate
which, if not exceeded, ensures that the associated queues are
short or empty. EF is intended to support services that offer
tightly-bounded loss, delay, and delay jitter.
Karn, et al. Best Current Practice [Page 27]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
3. The Assured Forwarding (AF) PHB group [RFC2597] offers different
levels of forwarding assurance for each aggregated flow of
packets. Each AF group is independently allocated forwarding
resources. Packets are marked with one of three drop
precedences; those with the highest drop precedence are dropped
with lower probability than those marked with the lowest drop
precedence. DSCPs are recommended for four independent AF
groups, although a DS domain can have more or fewer AF groups.
Ongoing work in the IETF is addressing ways to support Intserv with
Diffserv. There is some belief (e.g., as expressed in [RFC2990])
that such an approach will allow individual flows to receive service
guarantees and scale to the global Internet.
The QoS guarantees that can be offered by the IP layer are a product
of two factors:
1. the concatenation of the QoS guarantees offered by the subnets
along the path of a flow. This implies that a subnet may wish to
offer multiple services (with different QoS guarantees) to the IP
layer, which can then determine which flows use which subnet
service. To put it another way, forwarding behavior in the
subnet needs to be "clued" by the forwarding behavior (service or
PHB) at the IP layer, and
2. the operation of a set of cooperating mechanisms, such as
bandwidth reservation and admission control, policy management,
traffic classification, traffic conditioning (marking, policing
and/or shaping), selective discard, queuing, and scheduling.
Note that support for QoS in subnets may require similar
mechanisms, especially when these subnets are general topology
subnets (e.g., ATM, frame relay, or MPLS) or shared media
subnets.
Many subnetwork designers face inherent tradeoffs between delay,
throughput, reliability, and cost. Other subnetworks have parameters
that manage bandwidth, internal connection state, and the like.
Therefore, the following subnetwork capabilities may be desirable,
although some might be trivial or moot if the subnet is a dedicated
point-to-point link.
1. The subnetwork should have the ability to reserve bandwidth for a
connection or flow and schedule packets accordingly.
2. Bandwidth reservations should be based on a one- or two-token
bucket model, depending on whether the service is intended to
support constant-rate or bursty traffic.
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3. If a connection or flow does not use its reserved bandwidth at a
given time, the unused bandwidth should be available for other
flows.
4. Packets in excess of a connection or flow's agreed rate should be
forwarded as best-effort or discarded, depending on the service
offered by the subnet to the IP layer.
5. If a subnet contains error control mechanisms (retransmission
and/or FEC), it should be possible for the IP layer to influence
the inherent tradeoffs between uncorrected errors, packet losses,
and delay. These capabilities at the subnet/IP layer service
boundary correspond to selection of more or less error control
and/or to selection of particular error control mechanisms within
the subnetwork.
6. The subnet layer should know, and be able to inform the IP layer,
how much fixed delay and delay jitter it offers for a flow or
connection. If the Intserv model is used, the delay jitter
component may be best expressed in terms of the TSpec/RSpec model
described in [RFC2212].
7. Support of the Diffserv class selectors [RFC2474] suggests that
the subnet might consider mechanisms that support priorities.
Subnetwork designers should be aware of the tradeoffs between
fairness and efficiency inherent in many transmission scheduling
algorithms. For example, many local area networks use contention
protocols to resolve access to a shared transmission channel. These
protocols represent overhead. While limiting the amount of data that
a subnet node may transmit per contention cycle helps assure timely
access to the channel for each subnet node, it also increases
contention overhead per unit of data sent.
In some mobile radio networks, capacity is limited by interference,
which in turn depends on average transmitter power. Some receivers
may require considerably more transmitter power (generating more
interference and consuming more channel capacity) than others.
In each case, the scheduling algorithm designer must balance
competing objectives: providing a fair share of capacity to each
subnet node while maximizing the total capacity of the network. One
approach for balancing performance and fairness is outlined in
[ES00].
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The TCP sender bases its retransmission timeout (RTO) on measurements
of the round trip delay experienced by previous packets. This allows
TCP to adapt automatically to the very wide range of delays found on
the Internet. The recommended algorithms are described in [RFC2988].
Evaluations of TCP's retransmission timer can be found in [AP99] and
[LS00].
These algorithms model the delay along an Internet path as a
normally-distributed random variable with a slowly-varying mean and
standard deviation. TCP estimates these two parameters by
exponentially smoothing individual delay measurements, and it sets
the RTO to the estimated mean delay plus some fixed number of
standard deviations. (The algorithm actually uses mean deviation as
an approximation to standard deviation, because it is easier to
compute.)
The goal is to compute an RTO that is small enough to detect and
recover from packet losses while minimizing unnecessary ("spurious")
retransmissions when packets are unexpectedly delayed but not lost.
Although these goals conflict, the algorithm works well when the
delay variance along the Internet path is low, or the packet loss
rate is low.
If the path delay variance is high, TCP sets an RTO that is much
larger than the mean of the measured delays. If the packet loss rate
is low, the large RTO is of little consequence, as timeouts occur
only rarely. Conversely, if the path delay variance is low, then TCP
recovers quickly from lost packets; again, the algorithm works well.
However, when delay variance and the packet loss rate are both high,
these algorithms perform poorly, especially when the mean delay is
also high.
Because TCP uses returning acknowledgments as a "clock" to time the
transmission of additional data, excessively high delays (even if the
delay variance is low) also affect TCP's ability to fully utilize a
high-speed transmission pipe. It also slows the recovery of lost
packets, even when delay variance is small.
Subnetwork designers should therefore minimize all three parameters
(delay, delay variance, and packet loss) as much as possible.
In many subnetworks, these parameters are inherently in conflict.
For example, on a mobile radio channel, the subnetwork designer can
use retransmission (ARQ) and/or forward error correction (FEC) to
trade off delay, delay variance, and packet loss in an effort to
improve TCP performance. While ARQ increases delay variance, FEC
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does not. However, FEC (especially when combined with interleaving)
often increases mean delay, even on good channels where ARQ
retransmissions are not needed and ARQ would not increase either the
delay or the delay variance.
The tradeoffs among these error control mechanisms and their
interactions with TCP can be quite complex, and are the subject of
much ongoing research. We therefore recommend that subnetwork
designers provide as much flexibility as possible in the
implementation of these mechanisms, and provide access to them as
discussed above in the section on Quality of Service.
Some subnetworks may provide asymmetric bandwidth (or may cause TCP
packet flows to experience asymmetry in the capacity) and the
Internet protocol suite will generally still work fine. However,
there is a case when such a scenario reduces TCP performance. Since
TCP data segments are "clocked" out by returning acknowledgments, TCP
senders are limited by the rate at which ACKs can be returned
[BPK98]. Therefore, when the ratio of the available capacity of the
Internet path carrying the data to the bandwidth of the return path
of the acknowledgments is too large, the slow return of the ACKs
directly impacts performance. Since ACKs are generally smaller than
data segments, TCP can tolerate some asymmetry, but as a general
rule, designers of subnetworks should be aware that subnetworks with
significant asymmetry can result in reduced performance, unless
issues are taken to mitigate this [RFC3449].
Several strategies have been identified for reducing the impact of
asymmetry of the network path between two TCP end hosts, e.g.,
[RFC3449]. These techniques attempt to reduce the number of ACKs
transmitted over the return path (low bandwidth channel) by changes
at the end host(s), and/or by modification of subnetwork packet
forwarding. While these solutions may mitigate the performance
issues caused by asymmetric subnetworks, they do have associated cost
and may have other implications. A fuller discussion of strategies
and their implications is provided in [RFC3449].
Many subnets include multiple links with varying traffic demands and
possibly different transmission speeds. At each link there must be a
queuing system, including buffering, scheduling, and a capability to
discard excess subnet packets. These queues may also be part of a
subnet flow control or congestion control scheme.
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For the purpose of this discussion, we talk about packets without
regard to whether they refer to a complete IP packet or a subnetwork
frame. At each queue, a packet experiences a delay that depends on
competing traffic and the scheduling discipline, and is subjected to
a local discarding policy.
Some subnets may have flow or congestion control mechanisms in
addition to packet dropping. Such mechanisms can operate on
components in the subnet layer, such as schedulers, shapers, or
discarders, and can affect the operation of IP forwarders at the
edges of the subnet. However, with the exception of Explicit
Congestion Notification [RFC3168] (discussed below), IP has no way to
pass explicit congestion or flow control signals to TCP.
TCP traffic, especially aggregated TCP traffic, is bursty. As a
result, instantaneous queue depths can vary dramatically, even in
nominally stable networks. For optimal performance, packets should
be dropped in a controlled fashion, not just when buffer space is
unavailable. How much buffer space should be supplied is still a
matter of debate, but as a rule of thumb, each node should have
enough buffering to hold one link_bandwidth*link_delay product's
worth of data for each TCP connection sharing the link.
This is often difficult to estimate, since it depends on parameters
beyond the subnetwork's control or knowledge. Internet nodes
generally do not implement admission control policies, and cannot
limit the number of TCP connections that use them. In general, it is
wise to err in favor of too much buffering rather than too little.
It may also be useful for subnets to incorporate mechanisms that
measure propagation delays to assist in buffer sizing calculations.
There is a rough consensus in the research community that active
queue management is important to improving fairness, link
utilization, and throughput [RFC2309]. Although there are questions
and concerns about the effectiveness of active queue management
(e.g., [MBDL99]), it is widely considered an improvement over tail-
drop discard policies.
One form of active queue management is the Random Early Detection
(RED) algorithm [RED93], a family of related algorithms. In one
version of RED, an exponentially-weighted moving average of the queue
depth is maintained:
When this average queue depth is between a maximum threshold
max_th and a minimum threshold min_th, the probability of packets
that are dropped is proportional to the amount by which the
average queue depth exceeds min_th.
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When this average queue depth is equal to max_th, the drop
probability is equal to a configurable parameter max_p.
When this average queue depth is greater than max_th, packets are
always dropped.
Numerous variants on RED appear in the literature, and there are
other active queue management algorithms which claim various
advantages over RED [GM02].
With an active queue management algorithm, dropped packets become a
feedback signal to trigger more appropriate congestion behavior by
the TCPs in the end hosts. Randomization of dropping tends to break
up the observed tendency of TCP windows belonging to different TCP
connections to become synchronized by correlated drops, and it also
imposes a degree of fairness on those connections that implement TCP
congestion avoidance properly. Another important property of active
queue management algorithms is that they attempt to keep average
queue depths short while accommodating large short-term bursts.
Since TCP neither knows nor cares whether congestive packet loss
occurs at the IP layer or in a subnet, it may be advisable for
subnets that perform queuing and discarding to consider implementing
some form of active queue management. This is especially true if
large aggregates of TCP connections are likely to share the same
queue. However, active queue management may be less effective in the
case of many queues carrying smaller aggregates of TCP connections,
e.g., in an ATM switch that implements per-VC queuing.
Note that the performance of active queue management algorithms is
highly sensitive to settings of configurable parameters, and also to
factors such as RTT [MBB00] [FB00].
Some subnets, most notably ATM, perform segmentation and reassembly
at the subnetwork edges. Care should be taken here in designing
discard policies. If the subnet discards a fragment of an IP packet,
then the remaining fragments become an unproductive load on the
subnet that can markedly degrade end-to-end performance [RF95].
Subnetworks should therefore attempt to discard these extra fragments
whenever one of them must be discarded. If the IP packet has already
been partially forwarded when discarding becomes necessary, then
every remaining fragment except the one marking the end of the IP
packet should also be discarded. For ATM subnets, this specifically
means using Early Packet Discard and Partial Packet Discard [ATMFTM].
Some subnets include flow control mechanisms that effectively require
that the rate of traffic flows be shaped upon entry to the subnet.
One example of such a subnet mechanism is in the ATM Available Bit
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rate (ABR) service category [ATMFTM]. Such flow control mechanisms
have the effect of making the subnet nearly lossless by pushing
congestion into the IP routers at the edges of the subnet. In such a
case, adequate buffering and discard policies are needed in these
routers to deal with a subnet that appears to have varying bandwidth.
Whether there is a benefit in this kind of flow control is
controversial; there are numerous simulation and analytical studies
that go both ways. It appears that some of the issues leading to
such different results include sensitivity to ABR parameters, use of
binary rather than explicit rate feedback, use (or not) of per-VC
queuing, and the specific ATM switch algorithms selected for the
study. Anecdotally, some large networks that used IP over ABR to
carry TCP traffic have claimed it to be successful, but have
published no results.
Another possible approach to flow control in the subnet would be to
work with TCP Explicit Congestion Notification (ECN) semantics
[RFC3168] through utilizing explicit congestion indicators in subnet
frames. Routers at the edges of the subnet, rather than shaping,
would set the explicit congestion bit in those IP packets that are
received in subnet frames that have an ECN indication. Nodes in the
subnet would need to implement an active queue management protocol
that marks subnet frames instead of dropping them.
ECN is currently a proposed standard, but it is not yet widely
deployed.
Application data compression is a function that can usually be
omitted in the subnetwork. The endpoints typically have more CPU and
memory resources to run a compression algorithm and a better
understanding of what is being compressed. End-to-end compression
benefits every network element in the path, while subnetwork-layer
compression, by definition, benefits only a single subnetwork.
Data presented to the subnetwork layer may already be in a compressed
format (e.g., a JPEG file), compressed at the application layer
(e.g., the optional "gzip", "compress", and "deflate" compression in
HTTP/1.1 [RFC2616]), or compressed at the IP layer (the IP Payload
Compression Protocol [RFC3173] supports DEFLATE [RFC2394] and LZS
[RFC2395]). Compression at the subnetwork edges is of no benefit for
any of these cases.
The subnetwork may also process data that has been encrypted by the
application (OpenPGP [RFC2440] or S/MIME [RFC2633]), just above TCP
(SSL, TLS [RFC2246]), or just above IP (IPsec ESP [RFC2406]).
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RFC 3819 Advice for Internet Subnetwork Designers July 2004
Ciphers generate high-entropy bit streams lacking any patterns that
can be exploited by a compression algorithm.
However, much data is still transmitted uncompressed over the
Internet, so subnetwork compression may be beneficial. Any
subnetwork compression algorithm must not expand uncompressible data,
e.g., data that has already been compressed or encrypted.
We make a strong recommendation that subnetworks operating at low
speed or with small MTUs compress IP and transport-level headers (TCP
and UDP) using several header compression schemes developed within
the IETF [RFC3150]. An uncompressed 40-byte TCP/IP header takes
about 33 milliseconds to send at 9600 bps. "VJ" TCP/IP header
compression [RFC1144] compresses most headers to 3-5 bytes, reducing
transmission time to several milliseconds on dialup modem links.
This is especially beneficial for small, latency-sensitive packets in
interactive sessions.
Similarly, RTP compression schemes, such as CRTP [RFC2508] and ROHC
[RFC3095], compress most IP/UDP/RTP headers to 1-4 bytes. The
resulting savings are especially significant when audio packets are
kept small to minimize store-and-forward latency.
Designers should consider the effect of the subnetwork error rate on
the performance of header compression. TCP ordinarily recovers from
lost packets by retransmitting only those packets that were actually
lost; packets arriving correctly after a packet loss are kept on a
resequencing queue and do not need to be retransmitted. In VJ TCP/IP
[RFC1144] header compression, however, the receiver cannot explicitly
notify a sender of data corruption and subsequent loss of
synchronization between compressor and decompressor. It relies
instead on TCP retransmission to re-synchronize the decompressor.
After a packet is lost, the decompressor must discard every
subsequent packet, even if the subnetwork makes no further errors,
until the sending TCP retransmits to re-synchronize the decompressor.
This effect can substantially magnify the effect of subnetwork packet
losses if the sending TCP window is large, as it will often be on a
path with a large bandwidth*delay product [LRKOJ99].
Alternate header compression schemes, such as those described in
[RFC2507], include an explicit request for retransmission of an
uncompressed packet to allow decompressor resynchronization without
waiting for a TCP retransmission. However, these schemes are not yet
in widespread use.
Both TCP header compression schemes do not compress widely-used TCP
options such as selective acknowledgements (SACK). Both fail to
compress TCP traffic that makes use of explicit congestion
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notification (ECN). Work is under way in the IETF ROHC WG to address
these shortcomings in a ROHC header compression scheme for TCP
[RFC3095] [RFC3096].
The subnetwork error rate also is important for RTP header
compression. CRTP uses delta encoding, so a packet loss on the link
causes uncertainty about the subsequent packets, which often must be
discarded until the decompressor has notified the compressor and the
compressor has sent re-synchronizing information. This typically
takes slightly more than the end-to-end path round-trip time. For
links that combine significant error rates with latencies that
require multiple packets to be in flight at a time, this leads to
significant error propagation, i.e., subsequent losses caused by an
initial loss.
For links that are both high-latency (multiple packets in flight from
a typical RTP stream) and error-prone, RTP ROHC provides a more
robust way of RTP header compression, at a cost of higher complexity
at the compressor and decompressor. For example, within a talk
spurt, only extended losses of (depending on the mode chosen) 12-64
packets typically cause error propagation.
The Internet architecture does not guarantee that packets will arrive
in the same order in which they were originally transmitted;
transport protocols like TCP must take this into account.
However, reordering does come at a cost with TCP as it is currently
defined. Because TCP returns a cumulative acknowledgment (ACK)
indicating the last in-order segment that has arrived, out-of-order
segments cause a TCP receiver to transmit a duplicate acknowledgment.
When the TCP sender notices three duplicate acknowledgments, it
assumes that a segment was dropped by the network and uses the fast
retransmit algorithm [Jac90] [RFC2581] to resend the segment. In
addition, the congestion window is reduced by half, effectively
halving TCP's sending rate. If a subnetwork reorders segments
significantly such that three duplicate ACKs are generated, the TCP
sender needlessly reduces the congestion window and performance
suffers.
Packet reordering frequently occurs in parts of the Internet, and it
seems to be difficult or impossible to eliminate [BPS99]. For this
reason, research on improving TCP's behavior in the face of packet
reordering [LK00] [BA02] has begun.
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[BPS99] cites reasons why it may even be undesirable to eliminate
reordering. There are situations where average packet latency can be
reduced, link efficiency can be increased, and/or reliability can be
improved if reordering is permitted. Examples include certain high
speed switches within the Internet backbone and the parallel links
used over many Internet paths for load splitting and redundancy.
This suggests that subnetwork implementers should try to avoid packet
reordering whenever possible, but not if doing so compromises
efficiency, impairs reliability, or increases average packet delay.
Note that every header compression scheme currently standardized for
the Internet requires in-order packet delivery on the link between
compressor and decompressor. PPP is frequently used to carry
compressed TCP/IP packets; since it was originally designed for
point-to-point and dialup links, it is assumed to provide in-order
delivery. For this reason, subnetwork implementers who provide PPP
interfaces to VPNs and other more complex subnetworks, must also
maintain in-order delivery of PPP frames.
Internet users are increasingly mobile. Not only are many Internet
nodes laptop computers, but pocket organizers and mobile embedded
systems are also becoming nodes on the Internet. These nodes may
connect to many different access points on the Internet over time,
and they expect this to be largely transparent to their activities.
Except when they are not connected to the Internet at all, and for
performance differences when they are connected, they expect that
everything will "just work" regardless of their current Internet
attachment point or local subnetwork technology.
Changing a host's Internet attachment point involves one or more of
the following steps.
First, if use of the local subnetwork is restricted, the user's
credentials must be verified and access granted. There are many ways
to do this. A trivial example would be an "Internet cafe" that
grants physical access to the subnetwork for a fee. Subnetworks may
implement technical access controls of their own; one example is IEEE
802.11 Wireless Equivalent Privacy [IEEE80211]. It is common
practice for both cellular telephone and Internet service providers
(ISPs) to agree to serve one anothers' users; RADIUS [RFC2865] is the
standard method for ISPs to exchange authorization information.
Second, the host may have to be reconfigured with IP parameters
appropriate for the local subnetwork. This usually includes setting
an IP address, default router, and domain name system (DNS) servers.
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On multiple-access networks, the Dynamic Host Configuration Protocol
(DHCP) [RFC2131] is almost universally used for this purpose. On PPP
links, these functions are performed by the IP Control Protocol
(IPCP) [RFC1332].
Third, traffic destined for the mobile host must be routed to its
current location. This roaming function is the most common meaning
of the term "Internet mobility".
Internet mobility can be provided at any of several layers in the
Internet protocol stack, and there is ongoing debate as to which is
the most appropriate and efficient. Mobility is already a feature of
certain application layer protocols; the Post Office Protocol (POP)
[RFC1939] and the Internet Message Access Protocol (IMAP) [RFC3501]
were created specifically to provide mobility in the receipt of
electronic mail.
Mobility can also be provided at the IP layer [RFC3344]. This
mechanism provides greater transparency, viz., IP addresses that
remain fixed as the nodes move, but at the cost of potentially
significant network overhead and increased delay because of the sub-
optimal network routing and tunneling involved.
Some subnetworks may provide internal mobility, transparent to IP, as
a feature of their own internal routing mechanisms. To the extent
that these simplify routing at the IP layer, reduce the need for
mechanisms like Mobile IP, or exploit mechanisms unique to the
subnetwork, this is generally desirable. This is especially true
when the subnetwork covers a relatively small geographic area and the
users move rapidly between the attachment points within that area.
Examples of internal mobility schemes include Ethernet switching and
intra-system handoff in cellular telephony.
However, if the subnetwork is physically large and connects to other
parts of the Internet at multiple geographic points, care should be
taken to optimize the wide-area routing of packets between nodes on
the external Internet and nodes on the subnet. This is generally
done with "nearest exit" routing strategies. Because a given
subnetwork may be unaware of the actual physical location of a
destination on another subnetwork, it simply routes packets bound for
the other subnetwork to the nearest router between the two. This
implies some awareness of IP addressing and routing within the
subnetwork. The subnetwork may wish to use IP routing internally for
wide area routing and restrict subnetwork-specific routing to
constrained geographic areas where the effects of suboptimal routing
are minimized.
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Subnetworks connecting more than two systems must provide their own
internal Layer-2 forwarding mechanisms, either implicitly (e.g.,
broadcast) or explicitly (e.g., switched). Since routing is the
major function of the Internet layer, the question naturally arises
as to the interaction between routing at the Internet layer and
routing in the subnet, and proper division of function between the
two.
Layer-2 subnetworks can be point-to-point, connecting two systems, or
multipoint. Multipoint subnetworks can be broadcast (e.g., shared
media or emulated) or non-broadcast. Generally, IP considers
multipoint subnetworks as broadcast, with shared-medium Ethernet as
the canonical (and historical) example, and point-to-point
subnetworks as a degenerate case. Non-broadcast subnetworks may
require additional mechanisms, e.g., above IP at the routing layer
[RFC2328].
IP is ignorant of the topology of the subnetwork layer. In
particular, reconfiguration of subnetwork paths is not tracked by the
IP layer. IP is only affected by whether it can send/receive packets
sent to the remotely connected systems via the subnetwork interface
(i.e., the reachability from one router to another). IP further
considers that subnetworks are largely static -- that both their
membership and existence are stable at routing timescales (tens of
seconds); changes to these are considered re-provisioning, rather
than routing.
Routing functionality in a subnetwork is related to addressing in
that subnetwork. Resolution of addresses on subnetwork links is
required for forwarding IP packets across links (e.g., ARP for IPv4,
or ND for IPv6). There is unlikely to be direct interaction between
subnetwork routing and IP routing. Where broadcast is provided or
explicitly emulated, address resolution can be used directly; where
not provided, the link layer routing may interface to a protocol for
resolution, e.g., to the Next-Hop Resolution Protocol [RFC2322] to
provide context-dependent address resolution capabilities.
Subnetwork routing can either complement or compete with IP routing.
It complements IP when a subnetwork encapsulates its internal
routing, and where the effects of that routing are not visible at the
IP layer. However, if different paths in the subnetwork have
characteristics that affect IP routing, it can affect or even inhibit
the convergence of IP routing.
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Routing protocols generally consider Layer-2 subnetworks, i.e., with
subnet masks and no intermediate IP hops, to have uniform routing
metrics to all members. Routing can break when a link's
characteristics do not match the routing metric, in this case, e.g.,
when some member pairs have different path characteristics. Consider
a virtual Ethernet subnetwork that includes both nearby (sub-
millisecond latency) and remote (100's of milliseconds away) systems.
Presenting that group as a single subnetwork means that some routing
protocols will assume that all pairs have the same delay, and that
that delay is small. Because this is not the case, the routing
tables constructed may be suboptimal or may even fail to converge.
When a subnetwork is used for transit between a set of routers, it
conventionally provides the equivalent of a full mesh of point-to-
point links. Simplicity of the internal subnet structure can be used
(e.g., via NHRP [RFC2332]) to reduce the size of address resolution
tables, but routing exchanges will continue to reflect the full mesh
they emulate. In general, subnetworks should not be used as a
transit among a set of routers where routing protocols would break if
a full mesh of equivalent point-to-point links were used.
Some subnetworks have special features that allow the use of more
effective or responsive routing mechanisms that cannot be implemented
in IP because of its need for generality. One example is the self-
learning bridge algorithm widely used in Ethernet networks. Learning
bridges perform Layer-2 subnetwork forwarding, avoiding the need for
dynamic routing at each subnetwork hop. Another is the "handoff"
mechanism in cellular telephone networks, particularly the "soft
handoff" scheme in IS-95 CDMA.
Subnetworks that cover large geographic areas or include links of
widely-varying capabilities should be avoided. IP routing generally
considers all multipoint subnets equivalent to a local, shared-medium
link with uniform metrics between any pair of systems, and ignores
internal subnetwork topology. Where a subnetwork diverges from that
assumption, it is the obligation of subnetwork designers to provide
compensating mechanisms. Not doing so can affect the scalability and
convergence of IP routing, as noted above.
The subnetwork designer who decides to implement internal routing
should consider whether a custom routing algorithm is warranted, or
if an existing Internet routing algorithm or protocol may suffice.
The designer should consider whether this decision is to reduce the
address resolution table size (possible, but with additional protocol
support required), or is trying to reduce routing table complexity.
The latter may be better achieved by partitioning the subnetwork,
either physically or logically, and using network-layer protocols to
support partitioning (e.g., AS's in BGP). Protocols and routing
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algorithms can be notoriously subtle, complex, and difficult to
implement correctly. Much work can be avoided if existing protocols
or implementations can be readily reused.
Security has become a high priority in the design and operation of
the Internet. The Internet is vast, and countless organizations and
individuals own and operate its various components. A consensus has
emerged for what might be called a "security placement principle": a
security mechanism is most effective when it is placed as close as
possible to, and under the direct control of the owner of the asset
that it protects.
A corollary of this principle is that end-to-end security (e.g.,
confidentiality, authentication, integrity, and access control)
cannot be ensured with subnetwork security mechanisms. Not only are
end-to-end security mechanisms much more closely associated with the
end-user assets they protect, they are also much more comprehensive.
For example, end-to-end security mechanisms cover gaps that can
appear when otherwise good subnetwork mechanisms are concatenated.
This is an important application of the end-to-end principle [SRC81].
Several security mechanisms that can be used end-to-end have already
been deployed in the Internet and are enjoying increasing use. The
most important are the Secure Sockets Layer (SSL) [SSL2] [SSL3] and
TLS [RFC2246] primarily used to protect web commerce, Pretty Good
Privacy (PGP) [RFC1991] and S/MIME [RFCs-2630-2634], primarily used
to protect and authenticate email and software distributions, the
Secure Shell (SSH), used for secure remote access and file transfer,
and IPsec [RFC2401], a general purpose encryption and authentication
mechanism that sits just above IP and can be used by any IP
application. (IPsec can actually be used either on an end-to-end
basis or between security gateways that do not include either or both
end systems.)
Nonetheless, end-to-end security mechanisms are not used as widely as
might be desired. However, the group could not reach consensus on
whether subnetwork designers should be actively encouraged to
implement mechanisms to protect user data.
The clear consensus of the working group held that subnetwork
security mechanisms, especially when weak or incorrectly implemented
[BGW01], may actually be counterproductive. The argument is that
subnetwork security mechanisms can lull end users into a false sense
of security, diminish the incentive to deploy effective end-to-end
Karn, et al. Best Current Practice [Page 41]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
mechanisms, and encourage "risky" uses of the Internet that would not
be made if users understood the inherent limits of subnetwork
security mechanisms.
The other point of view encourages subnetwork security on the
principle that it is better than the default situation, which all too
often is no security at all. Users of especially vulnerable subnets
(such as consumers who have wireless home networks and/or shared
media Internet access) often have control over at most one endpoint
-- usually a client -- and therefore cannot enforce the use of end-
to-end mechanisms. However, subnet security can be entirely adequate
for protecting low-valued assets against the most likely threats. In
any event, subnet mechanisms do not preclude the use of end-to-end
mechanisms, which are typically used to protect highly-valued assets.
This viewpoint recognizes that many security policies implicitly
assume that the entire end-to-end path is composed of a series of
concatenated links that are nominally physically secured. That is,
these policies assume that all endpoints of all links are trusted,
and that access to the physical medium by attackers is difficult. To
meet the assumptions of such policies, explicit mechanisms are needed
for links (especially shared medium links) that lack physical
protection. This, for example, is the rationale that underlies Wired
Equivalent Privacy (WEP) in the IEEE 802.11 [IEEE80211] wireless LAN
standard, and the Baseline Privacy Interface in the DOCSIS [DOCSIS1]
[DOCSIS2] data over cable television networks standards.
We therefore recommend that subnetwork designers who choose to
implement security mechanisms to protect user data be as candid as
possible with the details of such security mechanisms and the
inherent limits of even the most secure mechanisms when implemented
in a subnetwork rather than on an end-to-end basis.
In keeping with the "placement principle", a clear consensus exists
for another subnetwork security role: the protection of the
subnetwork itself. Possible threats to subnetwork assets include
theft of service and denial of service; shared media subnets tend to
be especially vulnerable to such attacks. In some cases, mechanisms
that protect subnet assets can also improve (but cannot ensure) end-
to-end security.
One security service can be provided by the subnetwork that will aid
in the solution of an overall Internet problem: subnetwork security
should provide a mechanism to authenticate the source of a subnetwork
frame. This function is missing in some current protocols, e.g., the
use of ARP [RFC826] to associate an IPv4 address with a MAC address.
The IPv6 Neighbor Discovery (ND) [RFC2461] performs a similar
function.
Karn, et al. Best Current Practice [Page 42]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
There are well-known security flaws with this address resolution
mechanism [Wilbur89]. However, the inclusion of subnetwork frame
source authentication will permit a secure subnetwork address.
Another potential role for subnetwork security is to protect users
against traffic analysis, i.e., identifying the communicating parties
and determining their communication patterns and volumes even when
their actual contents are protected by strong end-to-end security
mechanisms. Lower-layer security can be more effective against
traffic analysis due to its inherent ability to aggregate the
communications of multiple parties sharing the same physical
facilities while obscuring higher-layer protocol information that
indicates specific end points, such as IP addresses and TCP/UDP port
numbers.
However, traffic analysis is a notoriously subtle and difficult
threat to understand and defeat, far more so than threats to
confidentiality and integrity. We therefore urge extreme care in the
design of subnetwork security mechanisms specifically intended to
thwart traffic analysis.
Subnetwork designers must keep in mind that design and implementation
for security is difficult [Schneier00]. [Schneier95] describes
protocols and algorithms which are considered well-understood and
believed to be sound.
Poor design process, subtle design errors and flawed implementation
can result in gaping vulnerabilities. In recent years, a number of
subnet standards have had problems exposed. The following are
examples of mistakes that have been made:
1. Use of weak and untested algorithms [Crypto9912] [BGW01]. For a
variety of reasons, algorithms were chosen which had subtle
flaws, making them vulnerable to a variety of attacks.
2. Use of "security by obscurity" [Schneier00] [Crypto9912]. One
common mistake is to assume that keeping cryptographic algorithms
secret makes them more secure. This is intuitive, but wrong.
Full public disclosure early in the design process attracts peer
review by knowledgeable cryptographers. Exposure of flaws by
this review far outweighs any imagined benefit from forcing
attackers to reverse engineer security algorithms.
3. Inclusion of trapdoors [Schneier00] [Crypto9912]. Trapdoors are
flaws surreptitiously left in an algorithm to allow it to be
broken. This might be done to recover lost keys or to permit
surreptitious access by governmental agencies. Trapdoors can be
discovered and exploited by malicious attackers.
Karn, et al. Best Current Practice [Page 43]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
4. Sending passwords or other identifying information as clear text.
For many years, analog cellular telephones could be cloned and
used to steal service. The cloners merely eavesdropped on the
registration protocols that exchanged everything in clear text.
5. Keys which are common to all systems on a subnet [BGW01].
6. Incorrect use of a sound mechanism. For example [BGW01], one
subnet standard includes an initialization vector which is poorly
designed and poorly specified. A determined attacker can easily
recover multiple ciphertexts encrypted with the same key stream
and perform statistical attacks to decipher them.
7. Identifying information sent in clear text that can be resolved
to an individual, identifiable device. This creates a
vulnerability to attacks targeted to that device (or its owner).
8. Inability to renew and revoke shared secret information.
9. Insufficient key length.
10. Failure to address "man-in-the-middle" attacks, e.g., with mutual
authentication.
11. Failure to provide a form of replay detection, e.g., to prevent a
receiver from accepting packets from an attacker that simply
resends previously captured network traffic.
12. Failure to provide integrity mechanisms when providing
confidentiality schemes [Bel98].
This list is by no means comprehensive. Design problems are
difficult to avoid, but expert review is generally invaluable in
avoiding problems.
In addition, well-designed security protocols can be compromised by
implementation defects. Examples of such defects include use of
predictable pseudo-random numbers [RFC1750], vulnerability to buffer
overflow attacks due to unsafe use of certain I/O system calls
[WFBA2000], and inadvertent exposure of secret data.
This document represents a consensus of the members of the IETF
Performance Implications of Link Characteristics (PILC) working
group.
Karn, et al. Best Current Practice [Page 44]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
This document would not have been possible without the contributions
of a great number of people in the Performance Implications of Link
Characteristics Working Group. In particular, the following people
provided major contributions of text, editing, and advice on this
document: Mark Allman provided the final editing to complete this
document. Carsten Bormann provided text on robust header
compression. Gorry Fairhurst provided text on broadcast and
multicast issues, routing, and many valuable comments on the entire
document. Aaron Falk provided text on bandwidth on demand. Dan
Grossman provided text on many facets of the document. Reiner Ludwig
provided thorough document review and text on TCP vs. Link-Layer
Retransmission. Jamshid Mahdavi provided text on TCP performance
calculations. Saverio Mascolo provided feedback on the document.
Gabriel Montenegro provided feedback on the document. Marie-Jose
Montpetit provided text on bandwidth on demand. Joe Touch provided
text on multicast, broadcast, and routing, and Lloyd Wood provided
many valuable comments on versions of the document.
References of the form RFCnnnn are Internet Request for Comments
(RFC) documents available online at www.rfc-editor.org.
[802.1D] Information Technology Telecommunications and
information exchange between systems Local and
metropolitan area networks, Common specifications Media
access control (MAC) bridges, IEEE 802.1D, 1998. ISO
15802-3.
[802.1p] IEEE, 802.1p, Standard for Local and Metropolitan Area
Networks - Supplement to Media Access Control (MAC)
Bridges: Traffic Class Expediting and Multicast.
[AP99] Allman, M. and V. Paxson, On Estimating End-to-End
Network Path Properties, In Proceedings of ACM SIGCOMM
99.
[AR02] Acar, G. and C. Rosenberg, Weighted Fair Bandwidth-on-
Demand (WFBoD) for Geo-Stationary Satellite Networks
with On-Board Processing, Computer Networks, 39(1),
2002.
[ATMFTM] The ATM Forum, "Traffic Management Specification,
Version 4.0", April 1996, document af-tm-0056.000.
http://www.atmforum.com/
Karn, et al. Best Current Practice [Page 45]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
[BA02] Blanton, E. and M. Allman, On Making TCP More Robust to
Packet Reordering. ACM Computer Communication Review,
32(1), January 2002.
[Bel98] Bellovin, S., "Cryptography and the Internet", in
Proceedings of CRYPTO '98, August 1998.
http://www.research.att.com/~smb/papers/inet-crypto.pdf
[BGW01] Borisov, N., Goldberg, I. and D. Wagner, "Intercepting
Mobile Communications: The Insecurity of 802.11," In
Proceedings of ACM MobiCom, July 2001.
[BPK98] Balakrishnan, H., Padmanabhan, V. and R. Katz. "The
Effects of Asymmetry on TCP Performance." ACM Mobile
Networks and Applications (MONET), 1998.
[BPS99] Bennet,, J.C.R., Partridge, C. and N. Shectman, "Packet
Reordering is Not Pathological Network Behavior",
IEEE/ACM Transactions on Networking, Vol. 7, No. 6,
December 1999.
[CGMP] Farinacci D., Tweedly A. and T. Speakman, "Cisco Group
Management Protocol (CGMP)", 1996/1997.
ftp://ftpeng.cisco.com/ipmulticast/specs/cgmp.txt
[Crypto9912] Schneier, B., "European Cellular Encryption Algorithms"
Crypto-Gram, December 15, 1999.
http://www.counterpane.com
[DIX82] Digital Equipment Corp, Intel Corp, Xerox Corp,
Ethernet Local Area Network Specification Version 2.0,
November 1982.
[DOCSIS1] Data-Over-Cable Service Interface Specifications, Radio
Frequency Interface Specification 1.0, SP-RFI-I05-
991105, November 1999, Cable Television Laboratories,
Inc.
[DOCSIS2] Data-Over-Cable Service Interface Specifications, Radio
Frequency Interface Specification 1.1, SP-RFIv1.1-I05-
000714, July 2000, Cable Television Laboratories, Inc.
[DOCSIS3] Lai, W.S., "DOCSIS-Based Cable Networks: Impact of
Large Data Packets on Upstream Capacity", 14th ITC
Specialists Seminar on Access Networks and Systems,
Barcelona, Spain, April 25-27, 2001.
Karn, et al. Best Current Practice [Page 46]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
[EN301192] ETSI, European Broadcasting Union, Digital Video
Broadcasting (DVB); DVB Specification for Data
Broadcasting, European Standard (Telecommunications
Series) EN 301 192 v1.2.1(1999-06).
[ES00] Eckhardt, D. and P. Steenkiste, "Effort-limited Fair
(ELF) Scheduling for Wireless Networks, Proceedings of
IEEE Infocom 2000.
[FB00] Firoiu V. and M. Borden, "A Study of Active Queue
Management for Congestion Control" to appear in Infocom
2000.
[GM02] Grieco1, L. and S. Mascolo, "TCP Westwood and Easy RED
to Improve Fairness in High-Speed Networks",
Proceedings of the 7th International Workshop on
Protocols for High-Speed Networks, April 2002.
[IEEE8023] IEEE 802.3 CSMA/CD Access Method.
http://standards.ieee.org/
[IEEE80211] IEEE 802.11 Wireless LAN standard.
http://standards.ieee.org/
[ISO3309] ISO/IEC 3309:1991(E), "Information Technology -
Telecommunications and information exchange between
systems - High-level data link control (HDLC)
procedures - Frame structure", International
Organization For Standardization, Fourth edition 1991-
06-01.
[ISO13818] ISO/IEC, ISO/IEC 13818-1:2000(E) Information
Technology - Generic coding of moving pictures and
associated audio information: Systems, Second edition,
2000-12-01 International Organization for
Standardization and International Electrotechnical
Commission.
[ITU-I363] ITU-T I.363.5 B-ISDN ATM Adaptation Layer Specification
Type AAL5, International Standards Organisation (ISO),
1996.
[Jac90] Jacobson, V., Modified TCP Congestion Avoidance
Algorithm. Email to the end2end-interest mailing list,
April 1990.
ftp://ftp.ee.lbl.gov/email/vanj.90apr30.txt
Karn, et al. Best Current Practice [Page 47]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
[KY02] Khafizov, F. and M. Yavuz, Running TCP Over IS-2000,
Proceedings of IEEE ICC, 2002.
[LK00] Ludwig, R. and R. H. Katz, "The Eifel Algorithm: Making
TCP Robust Against Spurious Retransmissions", ACM
Computer Communication Review, Vol. 30, No. 1, January
2000.
[LKJK02] Ludwig, R., Konrad, A., Joseph, A. D. and R. H. Katz,
"Optimizing the End-to-End Performance of Reliable
Flows over Wireless Links", Kluwer/ACM Wireless
Networks Journal, Vol. 8, Nos. 2/3, pp. 289-299,
March-May 2002.
[LRKOJ99] Ludwig, R., Rathonyi, B., Konrad, A., Oden, K. and A.
Joseph, Multi-Layer Tracing of TCP over a Reliable
Wireless Link, pp. 144-154, In Proceedings of ACM
SIGMETRICS 99.
[LS00] Ludwig, R. and K. Sklower, The Eifel Retransmission
Timer, ACM Computer Communication Review, Vol. 30, No.
3, July 2000.
[MAGMA-PROXY] Fenner, B., He, H., Haberman, B. and H. Sandick,
"IGMP/MLD-based Multicast Forwarding ("IGMP/MLD
Proxying")", Work in Progress.
[MAGMA-SNOOP] Christensen, M., Kimball, K. and F. Solensky,
"Considerations for IGMP and MLD Snooping Switches",
Work in Progress.
[MBB00] May, M., Bonald, T. and J-C. Bolot, "Analytic
Evaluation of RED Performance", INFOCOM 2000.
[MBDL99] May, M., Bolot, J., Diot, C. and B. Lyles, "Reasons not
to deploy RED", Proc. of 7th. International Workshop on
Quality of Service (IWQoS'99), June 1999.
[MSMO97] Mathis, M., Semke, J., Mahdavi, J. and T. Ott, "The
Macroscopic Behavior of the TCP Congestion Avoidance
Algorithm", Computer Communication Review, Vol. 27,
number 3, July 1997.
[MYR95] Boden, N., Cohen, D., Felderman, R., Kulawik, A.,
Seitz, C., et al. MYRINET: A Gigabit per Second Local
Area Network, IEEE-Micro, Vol. 15, No.1, February 1995,
pp. 29-36.
Karn, et al. Best Current Practice [Page 48]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
[PFTK98] Padhye, J., Firoiu, V., Towsley, D. and J. Kurose,
"Modeling TCP Throughput: a Simple Model and its
Empirical Validation", UMASS CMPSCI Tech Report TR98-
008, Feb. 1998.
[RED93] Floyd, S. and V. Jacobson, "Random Early Detection
gateways for Congestion Avoidance", IEEE/ACM
Transactions in Networking, Vol. 1 No. 4, August 1993.
http://www.aciri.org/floyd/papers/red/red.html
[RF95] Romanow, A. and S. Floyd, "Dynamics of TCP Traffic over
ATM Networks". IEEE Journal of Selected Areas in
Communication, Vol.13 No. 4, May 1995, p. 633-641.
[RFC791] Postel, J., "Internet Protocol", STD 5, RFC 791,
September 1981.
[RFC793] Postel, J., "Transmission Control Protocol", STD 7, RFC
793, September 1981.
[RFC768] Postel, J., "User Datagram Protocol", STD 6, RFC 768,
August 1980.
[RFC826] Plummer, D.C., "Ethernet Address Resolution Protocol:
Or converting network protocol addresses to 48-bit
Ethernet address for transmission on Ethernet
hardware", STD 37, RFC 826, November 1982.
[RFC1071] Braden, R., Borman, D. and C. Partridge, "Computing the
Internet checksum", RFC 1071, September 1988.
[RFC1112] Deering, S., "Host Extensions for IP Multicasting", STD
5, RFC 1112, August 1989.
[RFC1144] Jacobson, V., "Compressing TCP/IP Headers for Low-Speed
Serial Links", RFC 1144, February 1990.
[RFC1191] Mogul, J. and S. Deering, "Path MTU Discovery", RFC
1191, November 1990.
[RFC1332] McGregor, C., "The PPP Internet Protocol Control
Protocol (IPCP)", RFC 1332, May 1992.
[RFC1435] Knowles, S., "IESG Advice from Experience with Path MTU
Discovery", RFC 1435, March 1993.
Karn, et al. Best Current Practice [Page 49]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
[RFC1633] Braden, R., Clark, D. and S. Shenker, "Integrated
Services in the Internet Architecture: an Overview",
RFC 1633, June 1994.
[RFC1661] Simpson, W., "The Point-to-Point Protocol (PPP)", STD
51, RFC 1661, July 1994.
[RFC1662] Simpson, W., Ed., "PPP in HDLC-like Framing", STD 51,
RFC 1662, July 1994.
[RFC1750] Eastlake 3rd, D., Crocker, S. and J. Schiller,
"Randomness Recommendations for Security", RFC 1750,
December 1994.
[RFC1812] Baker, F., Ed., "Requirements for IP Version 4
Routers", RFC 1812, June 1995.
[RFC1939] Myers, J. and M. Rose, "Post Office Protocol - Version
3", STD 53, RFC 1939, May 1996.
[RFC1981] McCann, J., Deering, S. and J. Mogul, "Path MTU
Discovery for IP version 6", RFC 1981, August 1996.
[RFC1991] Atkins, D., Stallings, W. and P. Zimmermann, "PGP
Message Exchange Formats", RFC 1991, August 1996.
[RFC2018] Mathis, M., Mahdavi, J., Floyd, S. and A. Romanow, "TCP
Selective Acknowledgement Options", RFC 2018, October
1996.
[RFC2131] Droms, R., "Dynamic Host Configuration Protocol", RFC
2131, March 1997.
[RFC2205] Braden, R., Ed., Zhang, L., Berson, S., Herzog, S. and
S. Jamin, "Resource ReSerVation Protocol (RSVP) --
Version 1 Functional Specification", RFC 2205,
September 1997.
[RFC2208] Mankin, A., Baker, F., Braden, B., Bradner, S., O`Dell,
M., Romanow, A., Weinrib, A. and L. Zhang, "Resource
ReSerVation Protocol (RSVP) -- Version 1 Applicability
Statement Some Guidelines on Deployment", RFC 2208,
September 1997.
[RFC2210] Wroclawski, J., "The Use of RSVP with IETF Integrated
Services", RFC 2210, September 1997.
Karn, et al. Best Current Practice [Page 50]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
[RFC2211] Wroclawski, J., "Specification of the Controlled-Load
Network Element Service", RFC 2211, September 1997.
[RFC2212] Shenker, S., Partridge, C. and R. Guerin,
"Specification of Guaranteed Quality of Service", RFC
2212, September 1997.
[RFC2246] Dierks, T. and C. Allen, "The TLS Protocol Version
1.0", RFC 2246, January 1999.
[RFC2309] Braden, B., Clark, D., Crowcroft, J., Davie, B.,
Deering, S., Estrin, D., Floyd, S., Jacobson, V.,
Minshall, G., Partridge, C., Peterson, L.,
Ramakrishnan, K., Shenker, S., Wroclawski, J. and L.
Zhang, "Recommendations on Queue Management and
Congestion Avoidance in the Internet", RFC 2309, April
1998.
[RFC2322] van den Hout, K., Koopal, A. and R. van Mook,
"Management of IP numbers by peg-dhcp", RFC 2322, 1
April 1998.
[RFC2328] Moy, J., "OSPF Version 2", STD 54, RFC 2328, April
1998.
[RFC2332] Luciani, J., Katz, D., Piscitello, D., Cole, B. and N.
Doraswamy, "NBMA Next Hop Resolution Protocol (NHRP)",
RFC 2332, April 1998.
[RFC2364] Gross, G., Kaycee, M., Li, A., Malis, A. and J.
Stephens, "PPP Over AAL5", RFC 2364, July 1998.
[RFC2394] Pereira, R., "IP Payload Compression Using DEFLATE",
RFC 2394, December 1998.
[RFC2395] Friend, R. and R. Monsour, "IP Payload Compression
Using LZS", RFC 2395, December 1998.
[RFC2401] Kent, S. and R. Atkinson, "Security Architecture for
the Internet Protocol", RFC 2401, November 1998.
[RFC2406] Kent, S. and R. Atkinson, "IP Encapsulating Security
Payload (ESP)", RFC 2406, November 1998.
[RFC2440] Callas, J., Donnerhacke, L., Finney, H. and R. Thayer,
"OpenPGP Message Format", RFC 2440, November 1998.
Karn, et al. Best Current Practice [Page 51]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
[RFC2460] Deering, S. and R. Hinden, "Internet Protocol, Version
6 (IPv6) Specification", RFC 2460, December 1998.
[RFC2461] Narten, T., Nordmark, E. and W. Simpson, "Neighbor
Discovery for IP Version 6 (IPv6)", RFC 2461, December
1998.
[RFC2474] Nichols, K., Blake, S., Baker, F. and D. Black,
"Definition of the Differentiated Services Field (DS
Field) in the IPv4 and IPv6 Headers", RFC 2474,
December 1998.
[RFC2475] Blake, S., Black, D., Carlson, M., Davies, E., Wang, Z.
and W. Weiss, "An Architecture for Differentiated
Services", RFC 2475, December 1998.
[RFC2507] Degermark, M., Nordgren, B. and S. Pink, "IP Header
Compression", RFC 2507, February 1999.
[RFC2508] Casner, S. and V. Jacobson, "Compressing IP/UDP/RTP
Headers for Low-Speed Serial Links", RFC 2508, February
1999.
[RFC2581] Allman, M., Paxson, V. and W. Stevens, "TCP Congestion
Control", RFC 2581, April 1999.
[RFC2582] Floyd, S. and T. Henderson, "The NewReno Modification
to TCP's Fast Recovery Algorithm", RFC 2582, April
1999.
[RFC2597] Heinanen, J., Baker, F., Weiss, W. and J. Wroclawski,
"Assured Forwarding PHB Group", RFC 2597, June 1999.
[RFC2616] Fielding, R., Gettys, J., Mogul, J., Frystyk, H.,
Masinter, L., Leach, P. and T. Berners-Lee, "Hypertext
Transfer Protocol -- HTTP/1.1", RFC 2616, June 1999.
[RFC2630] Housley, R., "Cryptographic Message Syntax", RFC 2630,
June 1999.
[RFC2631] Rescorla, E., "Diffie-Hellman Key Agreement Method",
RFC 2631, June 1999.
[RFC2632] Ramsdell, B., Ed., "S/MIME Version 3 Certificate
Handling", RFC 2632, June 1999.
Karn, et al. Best Current Practice [Page 52]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
[RFC2633] Ramsdell, B., "S/MIME Version 3 Message Specification",
RFC 2633, June 1999.
[RFC2634] Hoffman, P., "Enhanced Security Services for S/MIME",
RFC 2634, June 1999.
[RFC2684] Grossman, D. and J. Heinanen, "Multiprotocol
Encapsulation over ATM Adaptation Layer 5", RFC 2684,
September 1999.
[RFC2686] Bormann, C., "The Multi-Class Extension to Multi-Link
PPP", RFC 2686, September 1999.
[RFC2687] Bormann, C., "PPP in a Real-time Oriented HDLC-like
Framing", RFC 2687, September 1999.
[RFC2689] Bormann, C., "Providing Integrated Services over Low-
bitrate Links", RFC 2689, September 1999.
[RFC2710] Deering, S., Fenner, W. and B. Haberman, "Multicast
Listener Discovery (MLD) for IPv6", RFC 2710, October
1999.
[RFC2784] Farinacci, D., Li, T., Hanks, S., Meyer, D. and P.
Traina, "Generic Routing Encapsulation (GRE)", RFC
2784, March 2000.
[RFC2865] Rigney, C., Willens, S., Rubens, A. and W. Simpson,
"Remote Authentication Dial In User Service (RADIUS)",
RFC 2865, June 2000.
[RFC2914] Floyd, S., "Congestion Control Principles", BCP 41, RFC
2914, September 2000.
[RFC2923] Lahey, K., "TCP Problems with Path MTU Discovery", RFC
2923, September 2000.
[RFC2988] Paxson, V. and M. Allman, "Computing TCP's
Retransmission Timer", RFC 2988, November 2000.
[RFC2990] Huston, G., "Next Steps for the IP QoS Architecture",
RFC 2990, November 2000.
[RFC3048] Whetten, B., Vicisano, L., Kermode, R., Handley, M.,
Floyd, S. and M. Luby, "Reliable Multicast Transport
Building Blocks for One-to-Many Bulk-Data Transfer",
RFC 3048, January 2001.
Karn, et al. Best Current Practice [Page 53]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
[RFC3095] Bormann, C., Ed., Burmeister, C., Degermark, M.,
Fukushima, H., Hannu, H., Jonsson, L-E., Hakenberg, R.,
Koren, T., Le, K., Liu, Z., Martensson, A., Miyazaki,
A., Svanbro, K., Wiebke, T., Yoshimura, T. and H.
Zheng, "RObust Header Compression (ROHC): Framework
and four profiles: RTP, UDP, ESP, and uncompressed",
RFC 3095, July 2001.
[RFC3096] Degermark, M., Ed., "Requirements for robust IP/UDP/RTP
header compression", RFC 3096, July 2001.
[RFC3150] Dawkins, S., Montenegro, G., Kojo, M. and V. Magret,
"End-to-end Performance Implications of Slow Links",
BCP 48, RFC 3150, July 2001.
[RFC3155] Dawkins, S., Montenegro, G., Kojo, M., Magret, V. and
N. Vaidya, "End-to-end Performance Implications of
Links with Errors", BCP 50, RFC 3155, August 2001.
[RFC3168] Ramakrishnan, K., Floyd, S. and D. Black, "The Addition
of Explicit Congestion Notification (ECN) to IP", RFC
3168, September 2001.
[RFC3173] Shacham, A., Monsour, B., Pereira, R. and M. Thomas,
"IP Payload Compression Protocol (IPComp)", RFC 3173,
September 2001.
[RFC3246] Davie, B., Charny, A., Bennet, J.C.R., Benson, K., Le
Boudec, J.Y., Courtney, W., Davari, S., Firoiu, V. and
D. Stiliadis, "An Expedited Forwarding PHB (Per-Hop
Behavior)", RFC 3246, March 2002.
[RFC3248] Armitage, G., Carpenter, B., Casati, A., Crowcroft, J.,
Halpern, J., Kumar, B. and J. Schnizlein, "A Delay
Bound alternative revision of RFC 2598", RFC 3248,
March 2002.
[RFC3344] Perkins, C., Ed., "IP Mobility Support for IPv4", RFC
3344, August 2002.
[RFC3366] Fairhurst, G. and L. Wood, "Advice to link designers on
link Automatic Repeat reQuest (ARQ)", BCP 62, RFC 3366,
August 2002.
Karn, et al. Best Current Practice [Page 54]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
[RFC3376] Cain, B., Deering, S., Kouvelas, I., Fenner, B. and A.
Thyagarajan, "Internet Group Management Protocol,
Version 3", RFC 3376, October 2002.
[RFC3449] Balakrishnan, H., Padmanabhan, V., Fairhurst, G. and M.
Sooriyabandara, "TCP Performance Implications of
Network Path Asymmetry", BCP 69, RFC 3449, December
2002.
[RFC3450] Luby, M., Gemmell, J., Vicisano, L., Rizzo, L. and J.
Crowcroft, "Asynchronous Layered Coding (ALC) Protocol
Instantiation", RFC 3450, December 2002.
[RFC3451] Luby, M., Gemmell, J., Vicisano, L., Rizzo, L.,
Handley, M. and J. Crowcroft, "Layered Coding Transport
(LCT) Building Block", RFC 3451, December 2002.
[RFC3452] Luby, M., Vicisano, L., Gemmell, J., Rizzo, L.,
Handley, M. and J. Crowcroft, "Forward Error Correction
(FEC) Building Block", RFC 3452, December 2002.
[RFC3453] Luby, M., Vicisano, L., Gemmell, J., Rizzo, L.,
Handley, M. and J. Crowcroft, "The Use of Forward Error
Correction (FEC) in Reliable Multicast", RFC 3453,
December 2002.
[RFC3488] Wu, I. and T. Eckert, "Cisco Systems Router-port Group
Management Protocol (RGMP)", RFC 3488, February 2003.
[RFC3501] Crispin, M., "INTERNET MESSAGE ACCESS PROTOCOL -
VERSION 4rev1", RFC 3501, March 2003.
[RFC3828] Larzon, L-A., Degermark, M., Pink, S., Jonsson, L-E.,
Ed. and G. Fairhurst, Ed., "The User Datagram Protocol
(UDP)-Lite Protocol", RFC 3828, June 2004.
[Schneier95] Schneier, B., Applied Cryptography: Protocols,
Algorithms and Source Code in C (John Wiley and Sons,
October 1995).
[Schneier00] Schneier, B., Secrets and Lies: Digital Security in a
Networked World (John Wiley and Sons, August 2000).
[SP2000] Stone, J. and C. Partridge, "When the CRC and TCP
Checksum Disagree", ACM SIGCOMM, September 2000.
http://www.acm.org/sigcomm/sigcomm2000/conf/
paper/sigcomm2000-9-1.pdf
Karn, et al. Best Current Practice [Page 55]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
[SRC81] Saltzer, J., Reed D. and D. Clark, "End-to-End
Arguments in System Design". Second International
Conference on Distributed Computing Systems (April,
1981) pages 509-512. Published with minor changes in
ACM Transactions in Computer Systems 2, 4, November,
1984, pages 277-288. Reprinted in Craig Partridge,
editor Innovations in internetworking. Artech House,
Norwood, MA, 1988, pages 195-206. ISBN 0-89006-337-0.
[SSL2] Hickman, K., "The SSL Protocol", Netscape
Communications Corp., Feb 9, 1995.
[SSL3] Frier, A., Karlton, P. and P. Kocher, "The SSL 3.0
Protocol", Netscape Communications Corp., Nov 18, 1996.
[TCPF98] Lin, D. and H.T. Kung, "TCP Fast Recovery Strategies:
Analysis and Improvements", IEEE Infocom, March 1998.
http://www.eecs.harvard.edu/networking/papers/infocom-
tcp-final-198.pdf
[WFBA2000] Wagner, D., Foster, J., Brewer, E. and A. Aiken, "A
First Step Toward Automated Detection of Buffer Overrun
Vulnerabilities", Proceedings of NDSS2000.
http://www.isoc.org/isoc/conferences/ndss/
2000/proceedings/039.pdf
[Wilbur89] Wilbur, Steve R., Jon Crowcroft, and Yuko Murayama.
"MAC layer Security Measures in Local Area Networks",
Local Area Network Security, Workshop LANSEC '89
Proceedings, Springer-Verlag, April 1989, pp. 53-64.
Karn, et al. Best Current Practice [Page 56]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
Aaron Falk
USC/Information Sciences Institute
4676 Admiralty Way
Marina Del Rey, CA 90292
Phone: 310-448-9327
EMail: falk@isi.edu
Saverio Mascolo
Dipartimento di Elettrotecnica ed Elettronica,
Politecnico di Bari Via Orabona 4, 70125 Bari, Italy
Phone: +39 080 596 3621
EMail: mascolo@poliba.it
URL: http://www-dee.poliba.it/dee-web/Personale/mascolo.html
Marie-Jose Montpetit
MJMontpetit.com
EMail: marie@mjmontpetit.com
Karn, et al. Best Current Practice [Page 57]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
Phil Karn, Editor
Qualcomm 5775 Morehouse Drive
San Diego CA 92121
Phone: 858 587 1121
EMail: karn@qualcomm.com
Carsten Bormann
Universitaet Bremen TZI
Postfach 330440
D-28334 Bremen, Germany
Phone: +49 421 218 7024
Fax: +49 421 218 7000
EMail: cabo@tzi.org
Godred (Gorry) Fairhurst
Department of Engineering, University of Aberdeen,
Aberdeen, AB24 3UE, United Kingdom
EMail: gorry@erg.abdn.ac.uk
URL: http://www.erg.abdn.ac.uk/users/gorry
Dan Grossman
Motorola, Inc.
111 Locke Drive
Marlboro, MA 01752
EMail: Dan.Grossman@motorola.com
Reiner Ludwig
Ericsson Research
Ericsson Allee
1 52134 Herzogenrath, Germany
Phone: +49 2407 575 719
EMail: Reiner.Ludwig@ericsson.com
Karn, et al. Best Current Practice [Page 58]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
Jamshid Mahdavi
Novell, Inc.
EMail: jmahdavi@earthlink.net
Gabriel Montenegro
Sun Microsystems Laboratories, Europe
180, Avenue de l'Europe
38334 Saint Ismier CEDEX
France
EMail: gab@sun.com
Joe Touch
USC/Information Sciences Institute
4676 Admiralty Way
Marina del Rey CA 90292
Phone: 310 448 9151
EMail: touch@isi.edu
URL: http://www.isi.edu/touch
Lloyd Wood
Cisco Systems
9 New Square Park, Bedfont Lakes
Feltham TW14 8HA
United Kingdom
Phone: +44 (0)20 8824 4236
EMail: lwood@cisco.com
URL: http://www.ee.surrey.ac.uk/Personal/L.Wood/
Karn, et al. Best Current Practice [Page 59]
RFC 3819 Advice for Internet Subnetwork Designers July 2004
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Karn, et al. Best Current Practice [Page 60]